Cryptographic computation method, cryptographic system, and computer program

ABSTRACT

A system and method for achieving secure and fast computation in hyperelliptic cryptography is realized. Fast scalar multiplication is achieve by executing computing operations including halving as computing processing in scalar multiplication with respect to a divisor D in hyperelliptic curve cryptography. For example, computing operations including halving are executed in scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x 2 +x+h 0 , f 4 =0 as parameters, a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x 2 +h 1 x+h 0 , f 4 =0 as parameters, or a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x as a parameter. Further, reduced complexity and faster computation are realized through the application of a table that records which of k 1 , k 1 ′, (k 0 , k 0 ′) is correct on the basis of a computed value of [½ i D] with respect to a fixed divisor D, and through a reduction in the number of inversion operations.

CROSS REFERENCE TO RELATED APPLICATIONS

The present application claims priority to Japanese Patent Document Nos. P2004-287166 filed on Sep. 30, 2004, P2005-015071 filed on Jan. 24, 2005, and P2005-119587 filed on Apr. 18, 2005, the contents of which are herein incorporated by reference.

TECHNICAL FIELD

The present invention relates to a cryptographic computation method, a cryptographic system, and a computer program. More particularly, the present invention relates to a cryptographic computation method, a cryptographic system, and a computer program, which realize faster scalar multiplication in hyperelliptic curve cryptography.

BACKGROUND ART

With the recent advances in network communication and electronic business transaction, it is becoming increasingly important to ensure the security of communications. One of the methods used to ensure security is cryptography. At present, communications are carried out by using a variety of cryptographic techniques.

For example, there has been put into practical use a system in which a cryptographic module is embedded in a small device such as an IC card, and data transmission/reception is performed between the IC card and a reader/writer as a data reading/writing device, thereby carrying out authentication processing, or encryption/decryption of the transmitted/received data.

IC cards executing cryptographic processing, for example, are widely used in a variety of gates such as the entrance gate of a train station or in shopping centers, and the demands for smaller size and faster processing speed are becoming increasingly stringent.

Cryptographic schemes are roughly divided into a common key scheme and a public key scheme. The common-key scheme is also referred to as symmetric cryptography. In the common key scheme, the sender and the recipient both own a common key. A typical application of the common key scheme is DES (Data Encryption Standard). The characteristic feature of a DES algorithm is that both encryption and decryption can be executed using substantially the same algorithm.

A scheme adopting a configuration in which, as opposed to the above-mentioned common key scheme, the key owned by the sender and that owned by the recipient are different is the public key scheme or asymmetric cryptography. Unlike common key cryptography in which a common key is used for encryption and decryption, public-key cryptography proves advantageous in terms of key management because only one specific person needs to own a secret key that must be kept in secret. In comparison to common key cryptography, however, public key cryptography involves lower data processing speed. As such, in general, the public key cryptography is frequently used for the distribution of secret keys, digital signatures, or other such applications involving low data volume. Typical known examples of public key cryptography includes RSA (Rivest-Shamir-Adleman) cryptography and ECC (Elliptic Curve Cryptography).

Elliptic curve cryptography uses an elliptic curve y²=x³+ax+b (where 4a³+27b²≠0) over a prime field, an elliptic curve y²+xy=x³+ax²+b (where b≠0) over two extension fields, or the like. A set including an infinity point (O) added to a point on each of these curves forms a finite group for the addition, and the infinity point O becomes the unit element thereof. In the following description, the addition of points in the finite group is represented by the operator +. The addition P+Q of two different points P, Q in the finite group is referred to as the “point addition”, and the addition P+P=2P of two points P in the finite group is referred to as the “point doubling”. An operation of adding the point P to itself k times, that is, an operation of finding a point P+P+ . . . +P=kP, is referred to as the “scalar multiplication of a point”.

As is commonly known, the scalar multiplication of a point can be computed with point addition and point doubling. The addition of points, the doubling of a point, and the scalar multiplication of a point in affine coordinates (x, y) or projective coordinates (X, Y, Z) on an elliptic curve over the prime field and an elliptic curve over two extension fields are described in IEEE P1363/D13 Standard Specifications for Public Key Cryptography.

An example of a scheme in which elliptic curve cryptography is generalized is HECC (Hyper-Elliptic Curve Cryptography) system proposed by Koblitz and Cantor. The hyperelliptic curve cryptography is described in Non-Patent Documents 1 and 2.

In elliptic curve cryptography, if P denotes a point on an elliptic curve defined over a finite field Fq, and Q denotes a point kP(kεZ), that is, a point obtained as a result of the scalar multiplication of the point P, the problem of finding k from Q can be solved as a discrete logarithmic problem. On the other hand, in hyperelliptic curve cryptography, if D, denotes be a divisor equal to a formal sum of points and D₂ denotes a divisor defined as a scalar multiplication kD₁, then the problem of finding k from D₂ can be treated as a discrete logarithmic problem in a Jacobian variety on a hyperelliptic curve as a public key cryptography problem.

In the case of a hyperelliptic curve, a value characterizing the curve is a genus g. Let q be equal to p^(n) (q=p^(n)) where p denotes a prime number and n denotes a positive integer. In this case, a hyperelliptic curve C defined over the finite field Fq as a curve of the genus g is defined by the following equation: y ² +h(x)y=f(x) where h(x), f(x)εFq[x], f(x) is the monic polynomial of degree 2g+1.

The opposite point −P to a point P=(x, y) on the hyperelliptic curve C is defined as −P=(x, y+h(x)). A point for which P=−P is referred to as a ramification point.

As is commonly known, assuming the same level of security as elliptic curve cryptography, the processing size (the number of bits) of a definition field of hyperelliptic curve cryptography can be reduced to 1/g times the processing size of a definition field of elliptic curve cryptography. The small processing size proves advantageous in terms of implementation, which is regarded as one of the benefits of hyperelliptic curve cryptography.

Next, the basic principles of hyperelliptic curve cryptography will be described. As described above, in hyperelliptic curve cryptography, the problem of finding k from D₂ can be treated as a discrete logarithmic problem in a Jacobian variety on the hyperelliptic curve and as such can be treated as a problem in public key cryptography where D₁ is a divisor equal to a formal sum of points, and D₂ is a divisor defined as a scalar multiplication kD₁.

In this case, a divisor is equal to a formal sum of points and can be expressed by the form:

$\begin{matrix} {D\; = {\sum\limits_{i\;}^{\;}\;{m_{i}P_{i}}}} & \left\lbrack {{Formula}\mspace{14mu} 1} \right\rbrack \end{matrix}$

Further, a semi reduced divisor can be expressed by the form:

$\begin{matrix} {{D\; = {{\sum\limits_{i\;}^{\;}\;{m_{i}P_{i}}} - {\left( {\sum\limits_{i\;}^{\;}\; m_{i}} \right)P_{\infty}}}},{m_{i} \geq 0}} & \left\lbrack {{Formula}\mspace{14mu} 2} \right\rbrack \end{matrix}$

However, for P_(i)=(x_(i), y_(i)) and I≠j, a relation P_(i)≠P_(j) holds true.

Σm_(i) in the above equation is referred to as the weight of the divisor D. Further, a semi reduced divisor having a weight not exceeding the genus g is referred to as a reduced divisor.

Using the polynomials U and VεFq[x], any semi reduced devisor D in a Jacobian variety on the hyperelliptic curve can be expressed as D=(U, V). This expression is referred to as a Mumford expression. The Mumford expression is described in, for example, Non-Patent Document 3. U=Π(x−x _(i))^(m) ^(i) V(x _(i))=y _(i) V(x)² +V(x)h(x)−f(x)≡0 mod U(x), deg V<deg U  [Formula 3]

By using the Mumford expression, any reduced divisor D for a genus 2 can be expressed by a set of polynomials each having elements over the finite field set in the coefficients of the polynomial and having an order not exceeding 2. That is, the reduced divisor can be expressed as (U,V)=(x ₂ +u ₁ x+u ₀ ,v ₁ x+v ₀), or (U,V)=(x+x ₀ ,y ₀).

Further, the zero element is expressed as (U,V)=(1,0)=O

Next, the scalar multiplication of a divisor used in hyperelliptic curve cryptography will be described. The scalar multiplication of a divisor can be carried out as a combination of the addition of the divisor, which is referred to as an addition algorithm, and the doubling of the divisor. Major addition algorithms will be described below.

The first proposed practical algorithm is a Cantor algorithm. The Cantor algorithm is described in, for example, Non-Patent Documents 1 and 2. This Cantor algorithm is applicable to a divisor on a hyperelliptic curve of any genus. However, the drawback of this Cantor algorithm is that in comparison to an elliptic curve algorithm, the algorithm is complicated and has high complexity.

Harley proposed an algorithm in which, by limiting the algorithm to hyperelliptic curves of genus 2, a case-by-case differentiation is made depending on the weight of a divisor, and optimization is performed for each individual case to achieve a reduction in complexity. Since then, this Harley algorithm has been the subject of recent extensive studies on the improvement and extension of computation algorithms in the HECC (Hyper-Elliptic Curve Cryptography).

(a) In accordance with the Harley algorithm, the definition field is used as a prime field and the Mumford expression is adopted as an expression of a divisor on a curve with a genus 2. Examples of studies aimed at reducing the complexity of this algorithm include those disclosed in Non-Patent Document 4, Non-Patent Document 5, Non-Patent Document 6, and the like.

(b) In addition, an example of processing in which the definition field is extended with respect to two extension fields is reported in each of Non-Patent Document 7 and Non-Patent Document 8.

(c) Further, Non-Patent Documents 11, 12, 6 and 13 disclose studies according to which a reduction in complexity is accomplished by using the Mumford expression to express a divisor and adopting the weighted coordinates.

Processing using the Harley algorithm will be described with reference to FIG. 1. It should be noted that the present invention relates to a hyperelliptic curve of genus 2 defined over a finite field of a characteristic 2. In the following description, it is assumed that the genus of the curve is 2, and the characteristic of the definition field is 2.

FIG. 1A is a diagram showing a processing example of the addition of divisors, D₁+D₂, where D₁ and D₂ each denote a divisor with a genus 2. It should be noted that divisors D₁ and D₂ are expressed as follows: D₁=(U₁,V₁), D₂=(U₂,V₂). First, a case-by-case differentiation is made depending on the weight values of the divisors. That is, depending on the values of the respective weights of [D₁+D₂], the processing is differentiated for the following cases:

(1) weight 2+weight 2

(2) weight 2+weight 1

(3) Exceptional Processing 1

Next, in the case of addition of a weight 2 to a weight 2 itself, that is, in the case (1): weight 2+weight 2, if the greatest common denominator gcd(U₁, U₂) for the two divisors D₁=(U₁,V₁) and D₂=(U₂,V₂) is 1 (gcd(U₁, U₂)=1), the two divisors D₁=(U₁,V₁) and D₂=(U₂,V₂) do not include a common point or points opposite to each other. In this case, addition processing according to

(1a) HarleyADD,

that is, the Harley algorithm is carried out. The processing of (1a) HarleyADD is processing referred to as Most Frequent Case disclosed in, for example, Non-Patent Document 7. The Most Frequent Case is a case occurring with a high probability in the addition processing to find the sum of D₁+D₂ of divisors for a genus 2.

The processing of (1a) HarleyADD occurs with a very high probability. The probability with which some other exceptional processing occurs is very low. If the conditions of the most frequent case are not satisfied, that is, if the greatest common denominator gcd(U₁, U₂) for the two divisors D₁=(U₁,V₁) and D₂=(U₂,V₂)=1 is not satisfied,

(1b) Exceptional Processing 2 is carried out.

Also for the case (2) of weight 2+weight 1, in the same way, it is checked as to whether or not gcd(U₁, U₂)=1. If gcd(U₁, U₂)=1 is satisfied, ExHarADD^(2+1→2)  (2a) is carried out, and if gcd(U₁, U₂)=1 is not satisfied,

(2b) Exceptional Processing 3 (2b) is carried out.

Exceptional processing 1 in the case (3) is carried out for cases of weight setting other than those of the cases (1) and (2) mentioned above.

It should be noted that the algorithm of the addition processing for a genus 2 described above is disclosed in detail in Non-Patent Document 8 (Table 1, 2).

The flow of doubling operation for a genus 2 is shown in FIG. 1(B). The doubling operation is processing represented as D+D=2D. Let D₁=(U₁,V₁) be the input, and D₂=[2]D₁ be the output.

As in the case of addition, different kinds of processing are carried out depending on the weight of the devisor D, as follows:

(4): weight 2

(5): weight 1

(6): weight 0

In the case (4) of weight 2, it is checked as to whether or not the divisor includes a ramification point. If no ramification point is included, the processing of (4a) HarleyDBL is carried out. If the divisor includes a ramification point, (4b) Exceptional Processing 6 is carried out. The algorithm of the HarleyDBL processing is disclosed as being the most frequent case in, for example, Non-Patent Document 7. The algorithm of the HarleyDBL processing is shown below.

$\begin{matrix} {{{{Algorithm}\mspace{14mu} 1\;{HarleyDBL}}{{{Input}\text{:}\mspace{11mu} D_{1}} = \left( {U_{1},V_{1}} \right)}{{{Output}\text{:}\mspace{14mu} D_{2}} = \left( {U_{2},V_{2}} \right)}}{{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},{{\gcd\left( {h,U_{1}} \right)} = 1}}\left. {1.\mspace{14mu} U_{1}^{\prime}}\leftarrow U_{1}^{2} \right.{\left. {2.\mspace{14mu} S}\leftarrow{{h^{- 1}\left( {f + {h\; V_{1}} + V_{1}^{2}} \right)}/U_{1}} \right.,{{mod}\; U_{1}}}\left. {3.\mspace{14mu} V_{1}^{\prime}}\leftarrow{{SU}_{1} + V_{1}} \right.\left. {4.\mspace{14mu} U_{2}^{\prime}}\leftarrow\;{\left( {f + {h\; V_{1}^{\prime}} + {V^{\prime}}_{1}^{2}} \right)/U^{\prime}} \right.\left. {5.\mspace{14mu} U_{2}}\leftarrow{{MakeMonic}\left( U_{2}^{\prime} \right)} \right.\left. {6.\mspace{14mu} V_{2}}\leftarrow{V_{1}^{\prime} + {h\;{mod}\; U_{2}}} \right.{{7.\mspace{14mu}{return}\mspace{14mu} D_{2}} = \left( {U_{2},V_{2}} \right)}} & \left\lbrack {{Formula}\mspace{14mu} 4} \right\rbrack \end{matrix}$

As will be described later, this processing occurs with a very high probability. The probability of the occurrence of some other exceptional processing is very low. As described above, if the conditions of the most frequent case are not satisfied, Exceptional Processing 6 is carried out.

In the case of weight 1 as well, it is checked whether or not gcd(U₁, U₂)=1, and the processing of (5a), ExHarDBL^(1+1→2), or Exceptional Processing 7 as the processing (5b) is carried out. The algorithm of ExHarDBL^(1+1→2) is disclosed in Non-Patent Document 8[4.12. (a)].

As described above, HarleyADD and HarleyDBL are referred to as the most frequent case. If a divisor is generated at random and is subjected to addition or doubling, the processing of HarleyADD or HarleyDBL occurs with a very high probability. It should be noted that description about HarleyADD and HarleyDBL becoming the most frequent case is given in, for example, Non-Patent Document 14.

In accordance with Non-Patent Document 14, the probability of the occurrence of processing other than the above-mentioned most frequent case is O(1/q) where q denotes the number of elements in the definition field. Since q^(g) is a large number with a required size of about 160 in secure cryptographic applications, in practice, it can be regarded that only HarleyADD or HarleyDBL occurs.

Thus, when the addition algorithm of HECC (Hyper-Elliptic Curve Cryptography) is implemented as cryptographic computation means such as an IC card by using the Harley algorithm or an improved algorithm thereof, it is often the case that only

HarleyADD, and

HarleyDBL

are implemented, and other types of complicated exceptional processing that has almost zero probability of occurrence are not executed. Examples of the method applied to exceptional processing in this case include execution of a Cantor algorithm that does not require case-by-case differentiation depending on the weight value. Since the larger the genus, the greater the load of the complicated exceptional processing, this implementation method is specially described in Non-Patent Documents 9 and 10.

Next, scalar multiplication of a divisor in the algorithm of HECC (Hyper-Elliptic Curve Cryptography) will be described. In the algorithm of the HECC (Hyper-Elliptic Curve Cryptography), scalar multiplication of a divisor is carried out as a combination of hyperelliptic addition and hyperelliptic doubling. The algorithm of the scalar multiplication will be described by taking the basic binary method and basic double-and-add-always method as examples.

As described above, in elliptic curve cryptography, assuming that P denotes a point on an elliptic curve defined over a finite field Fq, Q denotes a point kP(kεZ), that is, a point obtained as a result of the scalar multiplication of the point P, the problem of finding k from Q can be solved as a discrete logarithmic problem. On the other hand, in hyperelliptic curve cryptography, assuming that D₁ denotes a divisor equal to a formal sum of points and D₂ denotes a divisor defined as a scalar multiplication kD₁, the problem of finding k from D₂ can be treated as a discrete logarithmic problem in a Jacobian variety on the hyperelliptic curve as a public key cryptography problem.

In this case, the binary expression of a scalar value: d as a multiplier to be applied to a scalar multiplication (D=dD) is given as follows: d=(d ₁₋₁ , . . . , d ₀)

wherein d₁₋₁=1, and d₁₋₂, . . . , ₀=1 or 0.

As the algorithm for scalar multiplication, the computation algorithms of the basic binary method include the following:

binary (left-to-right) method; and

binary (right-to-left) method.

According to the binary (right-to-left) method, d is scanned from the least significant bits, and if d_(i)=1, [2^(i)]D is added. The algorithm (Algorithm 2) of the binary (right-to-left) method is shown below.

$\begin{matrix} {{{Algorithm}\mspace{14mu} 2\mspace{14mu}{binary}\mspace{14mu}\left( {{right}\text{-}{to}\text{-}{left}} \right)\mspace{14mu}{method}}{{Input}\mspace{14mu} D_{0}}{{{Output}\mspace{14mu} D} = {d\; D_{0}}}\left. T\leftarrow D_{0} \right.\left. D\leftarrow O \right.{{{for}{\mspace{11mu}\;}{i{from}}\mspace{14mu} 0\mspace{14mu}{to}\mspace{14mu} l} - 1}\left\{ \mspace{20mu}{{{if}\mspace{14mu} d_{i}} = {\left. {1\mspace{14mu}{then}\mspace{14mu} D}\leftarrow{D + T} \right.//\left. {{Addition}\mspace{14mu}{HarleyADD}\mspace{20mu} T}\leftarrow{2\; T\text{//}{Doubling}\mspace{14mu}{HarleyDBL}} \right.}} \right\}{{return}\mspace{14mu} D}} & \left\lbrack {{Formula}\mspace{14mu} 5} \right\rbrack \end{matrix}$

On the other hand, according to the binary (left-to-right) method, d is scanned from the most significant bits, D is doubled for every bit, and if d_(i)=1, a base point is added. The algorithm (Algorithm 3) of the binary (left-to-right) method is shown below.

$\begin{matrix} {{{Algorithm}\mspace{14mu} 3\mspace{14mu}{binary}\mspace{14mu}\left( {{left}\text{-}{to}\text{-}{right}} \right)\mspace{14mu}{method}}{{Input}\mspace{14mu} D_{0}}{{{Output}\mspace{14mu} D} = {d\; D_{0}}}\left. D\leftarrow D_{0} \right.{{{for}{\mspace{11mu}\;}i\mspace{14mu}{from}\mspace{20mu} l} - {2\mspace{14mu}{downto}\mspace{14mu} 0}}\left\{ \mspace{20mu}{\left. D\leftarrow{\lbrack 2\rbrack D\text{//}{Doubling}\mspace{14mu}{HarleyDBL}\mspace{20mu}{if}\mspace{14mu} d_{i}} \right. = \left. {1\mspace{14mu}{then}\mspace{14mu} D}\leftarrow{D + {D_{0}\text{//}{Addition}\mspace{14mu}{HarleyADD}}} \right.} \right\}{{return}\mspace{14mu} D}} & \left\lbrack {{Formula}\mspace{14mu} 6} \right\rbrack \end{matrix}$

Next, base-point generation processing will be described. When applying scalar multiplication to cryptography, divisors D₀ necessary for the inputs are divided into the following two types:

(1): a divisor determined in advance; and

(2): a divisor undeterminable in advance and generated at random.

In the case of type (1) of a divisor determined in advance, the input divisor is referred to as a base point.

A general algorithm for generating a base point is described as follows.

(a): g elements on a definition field Fq are selected at random and g points P_(i) (where i=1, . . . , g) on a hyperelliptic curve are generated.

-   -   (a1): The elements selected at random are each used as an x         coordinate x_(i) (where i=1 . . . g). Then, a y coordinate         corresponding to x_(i) is determined so that every point (x, y)         is positioned on the hyperelliptic curve.

(b): Let D₀=(U(x), V(x)) represent the divisor of the base point.

-   -   (b1): U(x)=(x−x₁)(x−x₂) . . . (x−x_(g))     -   (b2): Coefficients v_(i) of an equation         V(x)=v_(g-1)x^(g−1)+v_(g-2)x^(g−2)+ . . . +v₀ are determined. If         the generated points are all different from each other, for         example, the coefficients v_(i) can be found from an equation         V(x_(i))=y_(i). (c): The divisors generated in accordance with         the above algorithm are each a divisor with a weight of g.

If the computation of scalar multiplication is applied to cryptography, a divisor D₀ required for the input, that is, a base point is generated. If divisors determined in advance are applied to the generation of a base point, it is possible to find a divisor with a weight of g as a divisor usable as a base point by carrying out the processing (a) to (c) described above.

Further, with regard to elliptic curve cryptography, halving of a rational point has been proposed. For example, halving of a rational point in elliptic curve cryptography is disclosed in Non-patent Document 15, Patent Document 1, and Patent Document 2. In the disclosed processing, when computing the scalar multiplication of a rational point, instead of using addition and doubling, addition and halving are used.

Halving in elliptic curve cryptography can be computed generally faster than doubling. As a result, scalar multiplications using halving can be computed fast. Non-patent Document 16 reports that in the case of a software implementation with the [Intel PentiumIII 800 MHz] from Intel Corporation as a processor, with respect to a definition field Fq, q=2¹⁶³, halving is approximately 2.1 times faster than doubling, and with respect to a definition field Fq, q=2²³³, halving is approximately 2.6 times faster than doubling. Since hyperelliptic curve cryptography represents the generalization of elliptic curve cryptography, there may be cases where the operations used in elliptic curve cryptography can be extended to hyperelliptic curve cryptography. For example, Non-patent Documents 17 and 18 disclose a case where the Montgomery method, which realizes fast computation since a y-coordinate is not used for computation in elliptic curve cryptography, is extended to hyperelliptic curve cryptography. It is anticipated that if halving faster than doubling can be realized also in hyperelliptic curve cryptography, the scalar multiplication of a divisor, too, can be computed faster than in the related art. However, the use of such halving operation is not known in the related art. It should be noted that Non-patent Document 19 is an example of a published document presenting a fast computation technique using doublings.

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SUMMARY

As opposed to the ECC (Elliptic Curve Cryptography) algorithm which is now entering the commercialization phase, the HECC (Hyper-Elliptic Curve Cryptography) algorithm, which is an extended concept of the ECC (Elliptic Curve Cryptography) algorithm, is currently under study at the academic-society level as to the construction of fast algorithms and their implementation methods. Nevertheless, the computation time of the scalar multiplication based on the HECC (Hyper-Elliptic Curve Cryptography) algorithm is still only approaching to that of the ECC (Elliptic Curve Cryptography) algorithm, and a further increase in computation speed is being desired.

The present invention has been made in view of the above-mentioned circumstances, and accordingly it is an object of the present invention to provide a cryptographic computation method, a cryptographic system, and a computer program, which enable a reduction in the computation time of scalar multiplication in HECC (Hyper-Elliptic Curve Cryptography) to realize fast HECC (Hyper-Elliptic Curve Cryptography) processing.

It is another object of the present invention to provide a cryptographic computation method, a cryptographic system, and a computer program, which find algorithms, curve parameters, and definition fields that allow halving in elliptic curve cryptography to be extended to hyperelliptic curve cryptography to achieve fast computation, thereby realizing fast computing processing through computing processing to which having is applied to hyperelliptic curve cryptography.

According to a first aspect of the present invention, there is provided a cryptographic computation method for executing cryptographic computation based on hyperelliptic curve cryptography, including a computing step of executing computing operations including halving as computing processing, in computation of scalar multiplication with respect to a divisor D on a hyperelliptic curve.

Further, in an embodiment of the cryptographic computation method according to the present invention, the computing step is a step of executing computing operations including halving in scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having a random parameter.

Further, in an embodiment of the cryptographic computation method according to the present invention, the computing step is a step of executing computing operations including halving in scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x²+x+h₀, f₄=0 as parameters.

Further, in an embodiment of the cryptographic computation method according to the present invention, the computing step is a step of executing computing operations including halving in scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x²+h₁x+h₀, f₄=0 as parameters.

Further, in an embodiment of the cryptographic computation method according to the present invention, the computing step is a step of executing computing operations including halving in scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x as a parameter.

Further, in an embodiment of the cryptographic computation method according to the present invention, the cryptographic computation method further includes a table-lookup step of looking up a table that records which of k₁, k₁′, (k₀, k₀′) is correct on the basis of a computed value of [½^(i)D] with respect to a divisor D fixed in advance, and the computing step executes computing processing in which complexity of halving is reduced, by determination processing based on a lookup of the table.

Further, in an embodiment of the cryptographic computation method according to the present invention, the computing step includes a step of calculating a value of an inverse 1/k₁ by multiplication and addition processing without performing an inversion, by application of the following relational expression: 1/k ₁ =h ₂ +k ₁ u ₂₁, which is derived from a halving computation algorithm in which Input: D ₂=(U ₂ ,V ₂), and Output: D ₁=(U ₁ ,V ₁)=[½]D ₂,

where U_(i)(x)=x²+u_(i1)x+u_(i0), V_(i)(x)=v_(i1)x+v_(i0), gcd(h, U_(i))=1, i=1, 2.

Further, in an embodiment of the cryptographic computation method according to the present invention, the cryptographic computation method executes computation according to an algorithm having a setting for not applying 1/u₂₁ as an input value, in a halving computation algorithm in which Input: D ₂=(U ₂ ,V ₂); and Output: D ₁=(U ₁ ,V ₁)=[½]D ₂,

where U_(i)(x)=x²+u_(i1)x+u_(i0), V_(i)(x)=v_(i1)x+v_(i0), gcd(h, U_(i))=1, i=1, 2.

Further, in an embodiment of the cryptographic computation method according to the present invention, the cryptographic computation method is a computation method for executing scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x²+h₁x+h₀, f₄=0 as parameters, and the computing step includes the step of setting, as an input value, 1/h₁ ² that is a previously calculated value, and applying the previously calculated input value 1/h₁ ² without executing processing of calculating an inverse 1/h₁ ².

Further, according to a second aspect of the present invention, there is provided a cryptographic system which executes cryptographic computation based on hyperelliptic curve cryptography, including a computation executing section that executes computing operations including halving as computing processing, in computation of scalar multiplication with respect to a divisor D on a hyperelliptic curve.

Further, in an embodiment of the cryptographic system according to the present invention, the computation executing section is configured to execute computing operations including halving in scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having a random parameter.

Further, in an embodiment of the cryptographic system according to the present invention, the computation executing section is configured to execute computing operations including halving in scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x²+x+h₀, f₄=0 as parameters.

Further, in an embodiment of the cryptographic system according to the present invention, the computation executing section is configured to execute computing operations including halving in scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x²+h₁x+h₀, f₄=0 as parameters.

Further, in an embodiment of the cryptographic system according to the present invention, the computation executing section is configured to execute computing operations including halving in scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x as a parameter.

Further, in an embodiment of the cryptographic system according to the present invention, the cryptographic system further includes a storage section that stores a table recording which of k₁, k₁′, (k₀, k₀′) is correct on the basis of a computed value of [½^(i)D] with respect to a divisor D fixed in advance, and the computation executing section is configured to execute computing processing in which complexity of doubling is reduced, by determination processing based on a lookup of the table.

Further, in an embodiment of the cryptographic system according to the present invention, the computation executing section is configured to calculate a value of an inverse 1/k₁ by multiplication and addition processing without performing an inversion, by application of the following relational expression: 1/k ₁ =h ₂ +k ₁ u ₂₁, which is derived from a halving computation algorithm in which Input: D ₂=(U ₂ ,V ₂), and Output: D ₁=(U ₁ ,V ₁)=[½]D ₂,

where U_(i)(x)=x²+u_(i1)x+u_(i0), V_(i)(x)=v_(i1)x+v_(i0), gcd(h, U_(i))=1, i=1, 2.

Further, in an embodiment of the cryptographic system according to the present invention, the computation executing section is configured to execute a halving computation algorithm in which Input: D ₂=(U ₂ ,V ₂), and Output: D ₁=(U ₁ ,V ₁)=[½]D ₂,

where U_(i)(x)=x²+u_(i1)x+u_(i0), V_(i)(x)=v_(i1)x+v_(i0), gcd(h, U_(i))=1, i=1, 2, and to execute computation according to an algorithm having a setting for not applying 1/u₂₁ as an input value.

Further, in an embodiment of the cryptographic system according to the present invention, the computation executing section is configured to execute scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x²+h₁x+h₀, f₄=0 as parameters, and to execute computation to which, with 1/h₁ ² that is a previously calculated value being set as an input value, the previously calculated input value 1/h₁ ² is applied without executing processing of calculating an inverse 1/h₁ ².

Further, according to a third aspect of the present invention, there is provided a computer program for causing cryptographic computation based on hyperelliptic curve cryptography to be executed on a computer, including a computing step of executing computing operations including halving as computing processing, in computation of scalar multiplication with respect to a divisor D on a hyperelliptic curve.

It should be noted that the computer program according to the present invention is a computer program that can be provided with respect to a computer system capable of executing a variety of program codes via a storage medium or communication medium that is provided in a computer-readable format, for example, a recording medium such as a CD, FD, or MO, or via a communication medium such as a network. By providing such a program in a computer-readable format, processing corresponding to that program is realized on the computer system.

Other objects, features, and advantages of the present invention will become apparent from the following detailed description of embodiments of the present invention and the accompanying drawings. It should be noted that the term system as used in this specification refers to a logical assembly of a plurality of devices, and is not limited to one in which devices of respective configurations are located within the same casing.

According to the configuration of the present invention, halving on elliptic curve cryptography is extended to hyperelliptic curve cryptography to thereby realize fast computation. In the case of cryptographic computation employing computations on a divisor on a hyperelliptic curve, a computing operation that puts a large load on the processing is the scalar multiplication of a divisor. Hence, by realizing faster scalar multiplication by the processing according to the present invention as described above, a considerable improvement can be achieved in terms of the processing of hyperelliptic curve cryptography.

According to the configuration of the present invention, in scalar multiplication with respect to a divisor D in hyperelliptic curve cryptography, faster scalar multiplication can be realized by executing computing operations including halving as computing processing. For example, fast computation is realized by executing computing operations including halving in scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x²+x+h₀, f₄=0 as parameters, a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x²+h₁x+h₀, f₄=0 as parameters, or a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x as a parameter.

According to the configuration of the present invention, a further reduction in the complexity of scalar multiplication of a divisor and hence faster computation can be achieved through the application of a table that records which of k₁, k₁′, (k₀, k₀′) is correct on the basis of a computed value of [½^(i)D] with respect to a divisor D fixed in advance.

According to the configuration of the present invention, in scalar multiplication with respect to a divisor D in hyperelliptic curve cryptography, computing operations including halving are executed as computing processing, and an algorithm for reducing the number of inversion operations executed in the halving computation processing is applied, thereby making it possible to achieve a further reduction in the complexity of scalar multiplication of a divisor and hence faster computation.

Additional features and advantages of the present invention are described in, and will be apparent from, the following Detailed Description and the Figures.

BRIEF DESCRIPTION OF THE FIGURES

FIG. 1 is a diagram illustrating the algorithms for the processing of addition and doubling in scalar multiplication in hyperelliptic curve cryptography on a curve of genus 2.

FIG. 2 is a diagram illustrating how doubling operation is differentiated on a case-by-case basis in hyperelliptic curve cryptography on a curve of genus 2.

FIG. 3 is a flow chart illustrating the algorithm of HEC_HLV.

FIG. 4 is a flow chart illustrating the algorithm of HEC_HLV^(2→1+1).

FIG. 5 is a flow chart illustrating the algorithm of HEC_HLV^(2→2+2).

FIG. 6 is a flow chart illustrating the algorithm of HEC_HLV^(1→2+2).

FIG. 7 is a block diagram showing the functional configuration of a cryptographic system executing cryptographic computation according to the present invention.

FIG. 8 is a diagram showing an example of the configuration of an IC module as an example of a cryptographic processing executing device that executes the cryptographic computation according to the present invention.

DETAILED DESCRIPTION

A cryptographic system and a cryptographic computation method, and a computer program according to the present invention will be described below in detail with reference to the drawings.

The present invention provides a fast computing method with respect to HECC (Hyper-Elliptic Curve Cryptography) that represents the generalization of elliptic curve cryptography. As described above, in the case of a hyperelliptic curve, the value characterizing the curve is a genus g. It is assumed that p denotes a prime number, n denotes a positive integer, and q=p^(n). In this case, a hyperelliptic curve C defined over the finite field Fq as a curve of the genus g is expressed by the following equation: y ² +h(x)y=f(x), where h(x), f(x)εFq[x], f(x) is the monic polynomial of degree 2g+1.

An opposite point −P to a point P=(x, y) on the hyperelliptic curve C is defined as −P=(x, y+h(x)). A point for which P=−P is referred to as a ramification point.

As is commonly known, assuming the same level of security as elliptic curve cryptography, the processing size (or the number of bits) of a definition field of the hyperelliptic curve cryptography can be reduced to 1/g times the processing size of a definition field of an elliptic curve cryptography. The small processing size proves advantageous in terms of implementation, which is considered to be one of the benefits of hyperelliptic curve cryptography.

Next, the basic principles of hyperelliptic curve cryptography will be described. As described above, in hyperelliptic curve cryptography, the problem of finding k from D₂ can be treated as a discrete logarithmic problem in a Jacobian variety on the hyperelliptic curve and as such can be treated as a problem in public key cryptography where D₁ is a divisor equal to a formal sum of points, and D₂ is a divisor defined as a scalar multiplication kD₁.

In this case, a divisor is equal to a formal sum of points and can be expressed by the form:

$\begin{matrix} {D\; = {\sum\limits_{i\;}{m_{i}P_{i}}}} & \left\lbrack {{Formula}\mspace{14mu} 7} \right\rbrack \end{matrix}$

Further, a semi reduced divisor can be expressed by the form:

$\begin{matrix} {{D\; = {{\sum\limits_{i\;}^{\;}\;{m_{i}P_{i}}} - {\left( {\sum\limits_{i\;}^{\;}\; m_{i}} \right)P_{\infty}}}},{m_{i} \geq 0}} & \left\lbrack {{Formula}\mspace{14mu} 8} \right\rbrack \end{matrix}$

However, for P_(i)=(x_(i), y_(i)) and I≠j, a relation P_(i)≠P_(j) holds true.

Σm_(i) in the above equation is referred to as the weight of the divisor D. Further, a semi reduced divisor having a weight not exceeding the genus g is referred to as a reduced divisor.

Using the polynomials U and VεFq[x], any semi reduced devisor D in a Jacobian variety on the hyperelliptic curve can be expressed as D=(U, V). This expression is referred to as a Mumford expression. U=Π(x−x _(i))^(m) ^(i) V(x _(i))=y _(i) V(x)² +V(x)h(x)−f(x)≡0 mod U(x), deg V<deg U  [Formula 9]

By using the Mumford expression, any reduced divisor D for a genus 2 can be expressed by a set of polynomials each having elements over the finite field set in the coefficients of the polynomial and having an order not exceeding 2. That is, the reduced divisor can be expressed as (U,V)=(x ₂ +u ₁ x+u ₀ ,v ₁ x+v ₀), or (U,V)=(x+x ₀ ,y ₀).

Further, the zero element is expressed as (U,V)=(1,0)=O.

According to the present invention, halving operation in elliptic curve cryptography is extended to hyperelliptic curve cryptography, and algorithms, curve parameters, and definition fields that allow faster computation than doubling operation are found, thereby realizing computing processing to which halving operation, which is faster than doubling operation, is applied to hyperelliptic curve cryptography. In the following, the description of embodiments of the present invention is organized in two parts. First, techniques according to Processing Examples 1 to 6 below will be described in the first part.

Processing Example 1 Proposed Method A1

A method of computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with random parameters.

Processing Example 2 Proposed Method F1

A method of computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+h₁x+h₀, f₄=0.

Processing Example 3 Proposed Method B1

A method of computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+x+h₀, f₄=0.

Processing Example 4 Proposed Method E1

A method of computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with a parameter h(x)=x.

Processing Example 5 Proposed Method C1

When computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with random parameters, a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+h₁x+h₀, f₄=0, and a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+x+h₀, f₄=0, two candidates of the halved value arise. In this case, it is necessary to select the one with the correct value from the two candidates. When selecting the correct one, it is necessary to compute the trace, multiplication, and square root of a finite field. Which one of the two candidates is correct depends on the divisor D. Hence, if the divisor D is fixed, information as to which one of the two candidates is correct is retained in a table in advance, and this table is looked up when selecting the correct value, thereby omitting the above-mentioned extra computations.

Processing Example 6 Proposed Method D1

A method of computing the scalar multiplication of a divisor by using the method of computing the halving of a divisor as set forth in each of Processing Examples 1 to 5.

Further, the second part will be directed to the description of the following techniques that represent improvements over Processing Examples 1 to 3 and Processing Examples 5 and 6.

Processing Example 7 Proposed Method A2

A method of computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with random parameters.

Processing Example 8 Proposed Method F2

A method of computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+h₁x+h₀, f₄=0.

Processing Example 9 Proposed Method B2

A method of computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+x+h₀, f₄=0.

Processing Example 10 Proposed Method C2

When computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with random parameters, a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+h₁x+h₀, f₄=0, and a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+x+h₀, f₄=0, two candidates of the halved value arise. In this case, it is necessary to select the one with the correct value from the two candidates. When selecting the correct one, it is necessary to compute the trace, multiplication, and square root of a finite field. Which one of the two candidates is correct depends on the divisor D. Hence, if the divisor D is fixed, information as to which one of the two candidates is correct is retained in a table in advance, and this table is looked up when selecting the correct value, thereby omitting the above-mentioned extra computations.

Processing Example 11 Proposed Method D2

A method of computing the scalar multiplication of a divisor by using the method of computing the halving of a divisor as set forth in each of Processing Examples 7 to 10.

The respective processing examples mentioned above will be sequentially described below in detail.

Processing Example 1 Proposed Method A1

Processing Example 1 (Proposed Method A1) relates to a method of computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with random parameters.

Further, it is assumed that the order of the divisors used in the processing below is r. That is, the divisors have no ramification point. It is assumed that the input divisors are as follows: D ₂=(U ₂ ,V ₂); U ₂ =u ₂₂ x ² +u ₂₁ x+u ₂₀; and V ₂ =v ₂₁ x+v ₂₀, where:

u₂₂=1 if the weight of the divisor D₂ is 2; and

u₂₂=1, u₂₁=1, and v₂₁=0 if the weight of the divisor D₂ is 1.

Since no ramification point is included, as the halving operation, it suffices to consider four inversion operations of ExHarDBL^(1+1→2), ExHarDBL^(2+2→1), ExHarDBL^(2+2→2), and HarleyDBL. Those other than HarleyDBL represent exceptional cases.

Here, ExHarDBL^(2+2→2) represents a computation in the case where the weight of the input divisor is 2, and the weight of the output divisor is 1. Further, ExHarDBL^(2+2→2) represents a computation in the case where the weight of the input divisor is 2, and the coefficient of the first order term of U₂ satisfies u₂₁=0 and the weight of the output divisor is 2. However, while ExHarDBL^(2+2→2) can be computed by HarleyDBL, since halving as the inversion operation thereof becomes an exceptional case, ExHarDBL^(2+2→2) is herein treated as an exceptional case.

The halving operations corresponding to ExHarDBL^(1+1→2), ExHarDBL^(2+2→1), ExHarDBL^(2+2→2), and HarleyDBL mentioned above are defined as ExHEC_HLV^(2→1+1), ExHEC_HLV^(1→2+2), ExHEC_HLV^(2→2+2), and HEC_HLV, respectively.

When carrying out the halving operation of a divisor, first, a case-by-case differentiation is made depending on the input divisor as shown in FIG. 2. If the weight of the input divisor is 2, and the coefficient of the first order term of U₂ satisfies u₂₁≠0, the computation is performed by HEC_HLV. Further, if the weight of the input divisor is 2, and the coefficient of the term of U₂ satisfies u₂₁=0, the computation is performed by ExHEC_HLV^(2→2+2) or ExHEC_HLV^(2→1+1). Further, if the weight of the input divisor is 1, the computation is performed by ExHEC_HLV^(1→2+2). The algorithm of HEC_HLV will be described with reference to FIG. 3.

FIG. 3 is a flowchart showing the algorithm of HEC_HLV.

In step S101, the inputs are defined as follows: D ₂=(U ₂ ,V ₂); U ₂ =x ² +u ₂₁ x+u ₂₀; and V ₂ =v ₂₁ x+v ₂₀.

In step S102, the roots of k₁h₂+k₁ ²u₂₁+1=0, k₁, k₁ are found, and in step S103, c₁←f₃+h₂v₂₁+u₂₀+(f₄+u₂₁)u₂₁ is set. In step S104, it is determined whether or not k₁h₀+k₀h₁+k₀ ²u₂₁+c₁=0 has roots. If it does not have roots, k₁←k₁′ is set in step S105, and if it has roots, the process advances to step S106 where the roots of k₁h₀+k₀h₁+k₀ ²u₂₁+c₁=0, k₀, k₀′ are found.

Next, the process advances to step S107 where u₁₁ is computed, and in step S108, it is determined whether or not xh₂+x²u₁₁+1=0 has roots. If it does not have roots, k₀←k₀′ is set in step S109, and if it has roots, the process advances to step S110 where u₁₀ is computed. Further, in step S111, v₁₁, v₁₀ are calculated, and in step S12, by setting as: U ₁ ←x ₂ +u ₁₁ x+u ₁₀; and V ₁ ←v ₁₁ x+v ₁₀, in step S113, the output D ₁←(U ₁ ,V ₁) is obtained.

The halving operation of a divisor is realized by the reverse operation of the algorithm for performing the doubling operation of a divisor, that is, Algorithm 1 [Algorithm 1 Harley DBL] below.

$\begin{matrix} {{{{Algorithm}\mspace{14mu} 1\;{HarleyDBL}}{{{Input}\text{:}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)}{{{Output}\text{:}\mspace{14mu} D_{2}} = \left( {U_{2},V_{2}} \right)}}\;{{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},{{\gcd\left( {h,U_{1}} \right)} = \left. {1{1.\mspace{14mu} U_{1}^{\prime}}}\leftarrow U_{1}^{2} \right.}}\left. {2.\mspace{14mu} S}\leftarrow{{{h^{- 1}\left( {f + {h\; V_{1}} + V_{1}^{2}} \right)}/U_{1}}\;{mod}\; U_{1}} \right.{\left. {3.\mspace{14mu} V_{1}^{\prime}}\leftarrow{{SU}_{1} + {V_{1}{4.\mspace{14mu} U_{2}^{\prime}}}}\leftarrow\;{{\left( {f + {h\; V_{1}^{\prime}} + {V^{\prime}}_{1}^{2}} \right)/U^{\prime}}{5.\mspace{14mu} U_{2}}}\leftarrow{{{MakeMonic}\left( U_{2}^{\prime} \right)}{6.\mspace{14mu} V_{2}}}\leftarrow{V_{1}^{\prime} + {h\;{mod}\; U_{2}7.\mspace{14mu}{return}\mspace{14mu} D_{2}}} \right. = \left( {U_{2},V_{2}} \right)}} & \left\lbrack {{Formula}\mspace{14mu} 10} \right\rbrack \end{matrix}$

In step 6 of Algorithm 1, there is a unique polynomial: k(x)=k ₁ x+k ₀, which satisfies: V ₁ ′+h=(k ₁ x+k ₀)U ₂ +V ₂.

This is transformed as follows. V ₁ ′=h+(k _(i) x+k ₀)U ₂ +V ₂. Substituting this into the expression that appears in step 4, (f+hV ₁ ′+V ₁′²), results in the following: U ₂ ′U ₁ ′=f+h(kU ₂ +V ₂)+k ² U ₂ ² +V ₂ ²  Expression (1).

In the above expression, since (U₂,V₂) is known, from Expression (1), the relational expression between k and U₁′ can be obtained.

In this case, it is to be noted that U ₂ =k ₁ ² U ₂.

The above equation (1) is expanded and rearranged to yield the following:

$\begin{matrix} {\;{{U_{\; 1}}^{\;\prime} = {x^{\; 4} + {\left( {\left( {{k_{\; 1}h_{\; 2}} + {k_{\; 1}^{\; 2}u_{\; 21}} + 1} \right)/k_{\; 1}^{\; 2}} \right)x^{\; 3}} + {\left( {\left( {{k_{\; 1}h_{\; 1}} + {k_{\; 0}h_{\; 2}} + {k_{\; 1}^{\; 2}u_{\; 20}} + k_{\; 0}^{\; 2} + c_{\; 2}} \right)/k_{\; 1}^{\; 2}} \right)x^{\; 2}} + {\left( {\left( {{k_{\; 1}h_{\; 0}} + {k_{\; 0}h_{\; 1}} + {k_{\; 0}^{\; 2}u_{\; 21}} + c_{\; 1}} \right)/k_{\; 1}^{\; 2}} \right)x} + {\left( {{k_{\; 0}h_{\; 0}} + {k_{\; 0}^{\; 2}u_{\; 20}} + c_{\; 0}} \right)/{k_{\; 1}^{\; 2}.}}}}} & {{Expression}\mspace{14mu}(2)} \end{matrix}$

Here, c ₂ =f ₄ +u ₂₁, c ₁ =f ₃ +h ₂ v ₂₁ +u ₂₁ +c ₂ u ₂₁, and c ₀ =f ₂ +h ₂ v ₂₀ +h ₁ v ₂₁ +v ₂₁ ² +c ₂ u ₂₀ +c ₁ u ₂₁ are satisfied.

Further, from step 1, U ₁ ′=U ₁ ², that is, the following expression holds: U ₁ ′=x ⁴ +u ₁₁ ² x ² +u ₁₀ ²  Expression (3)

A relational expression is derived through comparison between the respective coefficients of Expressions (2) and (3) mentioned above, and halving operation can be computed by solving this relational expression. The algorithm prescribing the above-mentioned procedure is shown below as Algorithm 4 [Algorithm 4 Sketch HEC_HLV].

$\begin{matrix} {{{{{{Algorithm}\mspace{14mu} 4\mspace{14mu}{Sketch}{\mspace{11mu}\;}{HEC\_ HLV}}{{{Input}\text{:}\mspace{14mu} D_{2}} = \left( {U_{2},V_{2}} \right)}{{{Output}\text{:}\mspace{14mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}}}{{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},{{\gcd\left( {h,U_{1}} \right)} = 1},{i = 1},{2{1.\mspace{14mu}{reconstruct}\mspace{14mu} k_{0}}},k_{1}}\mspace{20mu}{\left. {1.1\mspace{20mu} V_{1}^{\prime}}\leftarrow{V_{2} + h + {kU}_{2}} \right.,{{k(x)} = {{k_{1}x} + k_{0}}}}\mspace{20mu}\left. {1.2\mspace{14mu} U_{1}^{\prime}}\leftarrow\;{\left( {f + {h\; V_{1}^{\prime}} + {V^{\prime}}_{1}^{2}} \right)/\left( {k_{1}^{2}U_{2}} \right)} \right.\mspace{20mu}{{1.3\mspace{14mu}{derive}\mspace{14mu} k_{0}},{{k_{1}\mspace{14mu}{from}\mspace{14mu}{{coeff}\left( {U_{1}^{\prime},3} \right)}} = 0},{{{coeff}\left( {U_{1}^{\prime},1} \right)} = 0}}{{2.\mspace{14mu}{compute}\mspace{14mu} u_{11}\mspace{14mu}{by}\mspace{14mu}{substituting}\mspace{14mu} k_{0}},{k_{1}\mspace{14mu}{in}\mspace{14mu}{{coeff}\left( {U_{1}^{\prime},2} \right)}{3.\mspace{14mu}{compute}\mspace{14mu} u_{10}\mspace{14mu}{by}\mspace{14mu}{substituting}\mspace{14mu} k_{0}}},\left. {k_{1}\mspace{14mu}{in}\mspace{14mu}{{coeff}\left( {U_{1}^{\prime},0} \right)}{4.\mspace{14mu} U_{1}}}\leftarrow{x^{2} + {u_{11}x} + {u_{10}{5.\mspace{14mu} V_{1}}}}\leftarrow{V_{2} + h + {{kU}_{2}{mod}\; U_{1}}} \right.}}{{6.\mspace{14mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)}}\mspace{11mu}} & \left\lbrack {{Formula}\mspace{14mu} 11} \right\rbrack \end{matrix}$

Specifically, the following relational expressions can be obtained.

[Formula 12] k ₁ h ₂ +k ₁ ² u ₂₁+1=0  Expression (4) k ₁ h ₀ +k ₀ h ₁ +k ₀ ² u ₂₁ +c ₁=0  Expression (5) u ₁₁=√{square root over (k ₁ h ₁ +k ₀ h ₂ +k ₁ ² u ₂₀ +k ₀ ² +c ₂)}/k ₁  Expression (6) u ₁₀=√{square root over (k ₀ h ₀ +k ₀ ² u ₂₀ +c ₀)}/k ₁  Expression (7)

It is necessary to compute the correct values of k₀, k₁ from these relational expressions. This can be computed using the lemma as described below.

[Lemma 1]

It is assumed that that h(x) is an irreducible polynomial. In this case, there is only one k₁ that satisfies the expressions (4) and (5). Further, the expression (5) has roots only for the correct k₁. Further, there is only one k₀ that allows the computation of the halved divisor D₁ in Algorithm 4. Further, the following expression: xh ₂ +x ² u ₁₁+1=0 has roots only for the correct k₀.

The above-mentioned Lemma 1 was applied to Algorithm 4. The detailed computation method of halving is shown as Algorithm 5 [Algorithm 5 Sketch HEC_HLV] below.

[Expression 13]

Algorithm  5  HEC_HLV Input:D₂ = (U₂, V₂) Output:D₁ = (U₁, V₁) = [1/2]D₂ U_(i)(x) = x² + u_(i 1)x + u_(i 0 ), V_(i)(x) − v_(i 1)x + v_(i 0), gcd  (h, U_(i)) = 1, i = 1, 2 1.  Solve  k₁h₂ + k₁²u₂₁ + 1 = 0 α ← h₂/u₂₁, γ ← 1/(u₂₁α²)/ * 1/(u₂₁α²) = u₂₁/h₂² * /k₁ ← H(γ)α, k₁^(′) ← k₁ + α 2.  Select  correct  k₁  by  solving  k₁h₀ + k₀h₁ + k₀²u₂₁ + c₁ = 0 c₂ ← f₄ + u₂₁, c₁ ← f₃ + h₂v₂₁ + u₂₀ + c₂u₂₁ c₀ ← f₂ + h₂v₂₀ + h₁v₂₁ + v₂₁² + c₂u₂₀ + c₁u₂₁ α ← h₁/u₂₁, γ ← (c₁ + k₁h₀)/(u₂₁α²) if  Tr(γ) = 1  then  k₁ ← k₁^(′), γ ← (c₁ + k₁h₀)/(u₂₁α²) k₀ ← H(γ)α, k₁^(′) ← k₁ + α 3.  Select  correct  k₀  by  checking  trace  of  xh₂ + x² u₁₁ + 1 = 0 ${\left. u_{11}\leftarrow{\sqrt{{k_{1}h_{1}} + {k_{0}h_{2}} + {k_{1}^{2}u_{20}} + k_{0}^{2} + c_{2}}/k_{1}} \right.,{\left. \gamma\leftarrow{{u_{11}/h_{2}^{2}}{if}\mspace{14mu}{{Tr}(\gamma)}} \right. = {1\mspace{14mu}{then}}}}\mspace{11mu}$ ${\mspace{20mu}\;}{\left. k_{0}\leftarrow k_{0}^{\prime} \right.,\left. u_{11}\leftarrow{\sqrt{{k_{1}h_{1}} + {k_{0}h_{2}} + {k_{1}^{2}u_{20}} + k_{0}^{2} + c_{2}}/k_{1}} \right.}$ 4.  Compute  U₁ $\left. u_{10}\leftarrow{\sqrt{{k_{0}h_{0}} + {k_{0}^{2}u_{20}} + c_{0}}/k_{1}} \right.$ 5.  Compute  V₁ = V₂ + h + kU₂mod U₁ w ← h₂ + k₁u₂₁ + k₀ + k₁u₁₁ v₁₁  ← v₂₁ + h₁ + k₁u₂₀ + k₀u₂₁ + u₁₀k₁ + u₁₁w v₁₀  ← v₂₀ + h₀ + k₀u₂₀ + w 6.  U₁(x) ← x² + u₁₁x + u₁₀, V₁(x) ← v₁₁x + v₁₀ 7.  return  D₁ = (U₁, V₁)

In Algorithm 5 [Algorithm 5] mentioned above, if k₁′, k₀′ are correct values (in other words, if k₁, k₀ are not correct values), the complexity of the algorithm is: 32M+5S+6I+3SR+2H+2T.

Here, M, S, SR, H, and T mean multiplication, squaring, inversion, square root operation, half-trace (operation to find the root of a quadratic equation), and trace (determination as to whether roots exist for a quadratic equation), respectively. The complexity becomes the largest if these k₁′, k₀′ are correct values.

Next, if k₁, k₀ are correct values (in other words, if k₁′, k₀′ are not correct values), the complexity becomes the smallest, so the complexity can be reduced by 2M in step 2, and the complexity can be reduced by 2M+1SR in step 3. That is, the complexity in this case is 28M+5S+6I+2SR+2H+2T, and the complexity becomes the smallest.

Next, if k₁, k₀′ are correct values (in other words, if k₁′, k₀ are not correct values), the complexity can be reduced by 2M+1SR in step 3. That is, the complexity in this case becomes: 30M+5S+6I+2SR+2H+2T.

Lastly, if k₁′, k₀ are correct values (in other words, if k₁, k₀′ are not correct values), the complexity can be reduced by 2M in step 2. That is, the complexity in this case becomes: 30M+5S+6I+3SR+2H+2T.

Upon checking the probabilities with which the above-mentioned four cases occur by computer experiment, it was confirmed that they occur with substantially the same ratio. In the description that follows, it is assumed that the probabilities with which the above-mentioned four cases occur are substantially equal. The averaging of the complexities in the above-mentioned four cases yields 30M+5S+6I+2.5SR+2H+2T.

Next, the exceptional cases:

ExHEC_HLV^(2→1+1);

ExHEC_HLV^(1→2+2); and

ExHEC_HLV^(2→2+2)

are considered. Since the probabilities with which these exceptional cases occur are so low as to be negligible, no evaluation on complexity will be made.

First, the algorithm of ExHEC_HLV^(2→1+1) will be described with reference to the flow of FIG. 4.

ExHEC_HLV^(2→1+1) is realized by a reverse operation of ExHarDBL^(1+1→2). Assuming that the input divisors for ExHarDBL^(1+1→2) are D ₁=(U ₁ ,V ₁), U ₁ =x+u ₁₀ , V ₁ =v ₁₀, the output divisors: D ₂=(U ₂ ,V ₂)=2D ₁ , U ₂ =x ² +u ₂₀ x, V ₂ =v ₂₁ x+v ₂₀, can be computed as follows: U ₂ =x ² +u ₂₀=(x+u ₁₀)², v ₁₂=(u ₁₀ ⁴ +f ₃ u ₁₀ ² +f ₁ +h ₁ v ₁₀)/h(u ₁₀), and v ₂₀ =v ₁₀ +v ₂₁ u ₁₀.

Using these relational expressions, ExHEC_HLV^(2→1+1) is computed.

Let the input divisors be

D ₂=(U ₂ ,V ₂), U ₂ =x ² +u ₂₀ x, V ₂ =v ₂₁ x+v ₂₀ (flow of FIG. 4; step S201).

To obtain the output devisors D ₁=(U ₁ ,V ₁)=[½]D ₂ , U ₁ =x+u ₁₀ , V ₁ =v ₁₀, in step S202, let u₁₀=√u₂₀, and in step S203, let v₁₀=(v₂₁(u₁₀)+u₁₀ ⁴+f₃u₁₀ ²+f₁)/h₁, and in step S204, let U ₁ =x+u ₁₀ , V ₁ =v ₁₀. Then, in step S205, the output divisor D ₁=(U ₁ ,V ₁) is obtained.

Next, the processing procedure for ExHEC_HLV^(2→2+2) will be described with reference to the flow of FIG. 5. In step S301, the input divisors are assumed to be D ₂=(U ₂ ,V ₂), U ₂ =x ² +u ₂₀ , V ₂ =v ₂₁ x+v ₂₀.

In step S302, k₁h₂+1=0 is solved with respect to k₁ to yield k₁←1/h₂.

In step S303, let c ₂ ←f ₄, and c ₁ ←f ₃ +h ₂ v ₂₁ +u ₂₀ +u ₂₁ c ₂, and in step S304,

k₁h₀+k₀h₁=0 is solved with respect to k₀ to give k₀←(k₁h₀+c₁)/h₁.

Next, in step S305, u₁₁ is computed, and in step S306, it is determined whether or not xh ₂ +x ² u ₁₁+1=0 has roots. If it does not has roots, in step S307, the output D₁ is determined (step S308) by D ₁←HEC_HLV^(2→1)(D ₂)

On the other hand, if xh ₂ +x ² u ₁₁+1=0 has roots, the process advances to step S309 where u₁₀ is computed, and further in step S310, v₁₁, v₁₀ are computed. Then, in step S311, let U ₁ ←x ² +u ₁₁ x+u ₁₀, and V ₁ ←v ₁₁ x+v ₁₀, and in step S312, the output D ₁←(U ₁ ,V ₁) is obtained.

The processing of ExHEC_HLV^(2→2+2) is specifically carried out by the following procedure.

Supposing the input divisors are D ₂=(U ₂ ,V ₂), U ₂ =x ² +u ₂₀ , V ₂ =v ₂₁ x+v ₂₀, if U ₂ =x ² +u ₂₀, that is, if the first order term of U₂ is 0, there are two output divisor candidates, which are represented as D ₁=(x+√u ₂₀ ,V ₂(√u ₂₀)), and D ₁′=(x ₂ +u ₁₁ x+u ₁₀ ,v ₁₁ x+v ₁₀).

If D₁ is correct, the computation is carried out using ExHEC_HLV^(2→1+1).

If D₁ is correct, the computation is carried out using ExHEC_HLV^(2→2+2).

The determination as to which of the algorithms is to be used is made on the basis of the following procedure.

1. assume that D₁′ is correct.

2. Compute u₁₁.

3. Compute the trace Tr(h₂/u₁₁ ²) of xh₂+x²u₁₁+1=0. If Tr(h₂/u₁₁ ²)=0, then D₁′ is correct, so computation is carried out using ExHEC_HLV^(2→2+2). Otherwise, that is, if Tr(h₂/u₁₁ ²)=1, then D₁ is correct, so computation is carried out using ExHEC_HLV^(2→1+1).

The computation algorithm of ExHEC_HLV^(2→2+2) is shown below as Algorithm 6 [Algorithm 6].

$\begin{matrix} {{{{Algorithm}\mspace{14mu} 6\mspace{14mu}{ExHEC\_ HLV}^{2\rightarrow{2 + 2}}}{{Input}\text{:}\mspace{14mu} D_{2}} = {\left( {U_{2},V_{2}} \right) = \left( {{x^{2} + u_{20}},{{v_{21}x} + v_{20}}} \right)}}{{{{Output}\text{:}\mspace{14mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left( {{x^{2} + {u_{11}x} + u_{10}},{{v_{11}x} + v_{10}}} \right) = {{\left\lbrack {1/2} \right\rbrack D_{2}{\gcd\left( {h,U_{i}} \right)}} = 1}}}},{i = 1},2}{{{1.\mspace{14mu}{Solve}\mspace{14mu} k_{1}h_{2}} + 1} = 0}\;\left. k_{1}\leftarrow{1/h_{2}} \right.{{{2.\mspace{14mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{1}\mspace{14mu}{by}\mspace{11mu}{solving}\mspace{14mu} k_{1}h_{0}} + {k_{0}h_{1}} + k_{0}^{2} + c_{1}} = 0}{\left. c_{2}\leftarrow f_{4} \right.,\left. c_{1}\leftarrow{f_{3} + {h_{2}v_{21}} + u_{20} + {c_{2}u_{21}}} \right.}\left. c_{0}\leftarrow{f_{2} + {h_{2}v_{20}} + {\left( {h_{1} + v_{21}} \right)v_{21}} + {c_{2}u_{20}} + {c_{1}u_{21}}} \right.\left. k_{0}\leftarrow{\left( {{k_{1}h_{0}} + c_{1}} \right)/h_{1}} \right.{{{3.\mspace{14mu}{Select}\mspace{14mu}{correct}\mspace{14mu}{algorithm}\mspace{14mu}{by}\mspace{14mu}{checking}\mspace{14mu}{trace}\mspace{14mu}{of}\mspace{14mu}{xh}_{2}} + {x^{2}u_{11}} + 1} = 0}{\left. u_{11}\leftarrow{\sqrt{{k_{1}h_{1}} + {k_{0}h_{2}} + {k_{1}^{2}u_{20}} + k_{0}^{2} + c_{2}}/k_{1}} \right.,\left. \gamma\leftarrow{u_{11}/h_{2}^{2}} \right.}{{{if}\mspace{14mu}{{Tr}(\gamma)}} = {1\mspace{14mu}{then}}}{\left. D_{1}\leftarrow{{ExHEC\_ HLV}^{2\rightarrow{l + 1}}\left( D_{2} \right)} \right.,{{return}\mspace{14mu} D_{1}}}{4.\mspace{14mu}{Compute}\mspace{14mu} U_{1}}\left. u_{10}\leftarrow{\sqrt{{k_{0}h_{0}} + {k_{0}^{2}u_{20}} + c_{0}}/k_{1}} \right.{{5.\mspace{14mu}{Compute}\mspace{14mu} V_{1}} = {V_{2} + h + {{kU}_{2}\mspace{14mu}{mod}\; U_{1}}}}\left. w\leftarrow{h_{2} + k_{1}} \right.\left. v_{11}\leftarrow{h_{1} + {k_{1}u_{20}} + k_{0} + {u_{11}w}} \right.\left. v_{10}\leftarrow{v_{20} + h_{0} + {k_{0}u_{20}} + {u_{10}w}} \right.{\left. {6.\mspace{14mu}{U_{1}(x)}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,\left. {V_{1}(x)}\leftarrow{{v_{11}x} + v_{10}} \right.}{{7.\mspace{14mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)}} & \left\lbrack {{Formula}\mspace{14mu} 14} \right\rbrack \end{matrix}$

Next, the processing procedure for ExHEC_HLV^(1→2+2) will be described with reference to the flow of FIG. 6. In step S401, the input divisors are assumed to be D ₂=(U ₂ ,V ₂); U ₂ =x+u ₂₀; and V ₂ =v ₂₀.

In step S402, let c₃←(f₄+u₂₀, and in step S403, the root of k₁h₂+k₁ ²u₂₁+c₃=0, k₁, k₁′ are found, and in step S404, let c ₁ ←f ₃ +c ₃ u ₂₀, and c ₁ ←f ₂ +h ₂ v ₂₀ +c ₂ u ₂₀, and in step S405, it is determined whether or not k₁h₀+k₀h₁+k₀ ²+c₁=0 has roots. If it does not have roots, the process advances to step S407 after k₁←k₁′ is set in step S406, and if it has roots, the process advances to step S407 as it is.

In step S407, the roots of k₁h₀+k₀h₁+k₀ ²+c₁=0, k₀, k₀′ are found. Then, the process advances to step S408 where u₁₁ is computed, and in step S409, it is determined whether or not xh₂+x²u₁₁+1=0 has roots. If it does not have roots, the process advances to step S411 after k₀←k₀′ is set in step S410, and if it has roots, the process advances to step S411 as it is, and u₁₀ is computed. Further, in step S412, v₁₁, v₁₀ are computed. In step S413, let U ₁ ←x ² +u ₁₁ x+u ₁₀, and V ₁ ←v ₁₁ x+v ₁₀, and in step S414, the output D ₁←(U ₁ ,V ₁) is obtained.

While the computation procedure for ExHEC_HLV^(1→2+2) is similar to that for HEC_HLV, a large difference resides in the weight of the input divisor. f+hV₁′+V₁′² of ExHEC_HLV^(1→2+2) thus becomes a quintic monic polynomial. Hence, unlike in the case of HEC_HLV, U ₁′←(f+hV ₁ ′+V ₁′²)/U ₂ is not divided by k₁ ². The computation algorithm of ExHEC_HLV^(1→2+2) is shown below as [Algorithm 7].

$\begin{matrix} {{{{Algorithm}\mspace{14mu} 7\mspace{14mu}{ExHEC\_ HLV}^{1\rightarrow{2 + 2}}}{{{Input}\text{:}\mspace{14mu} D_{2}} = {\left( {U_{2},V_{2}} \right) = \left( {{x + u_{20}},v_{20}} \right)}}{{{Output}\text{:}\mspace{11mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left( {{x^{2} + {u_{11}x} + u_{10}},{{v_{11}x} + v_{10}}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}}}{{{\gcd\left( {h,U_{i}} \right)} = 1},{i = 1},2}{{{1.\mspace{14mu}{Solve}\mspace{14mu} k_{1}h_{2}} + {k_{1}^{2}u_{21}} + c_{3}} = 0}{\left. c_{3}\leftarrow{f_{4} + u_{20}} \right.,\left. \alpha\leftarrow h_{2} \right.,\left. \gamma\leftarrow{c_{3}/\left( {u_{21}\alpha^{2}} \right)} \right.}{\left. k_{1}\leftarrow{{H(\gamma)}\alpha} \right.,\left. k_{1}^{\prime}\leftarrow{k_{1} + \alpha} \right.}{{{2.\mspace{14mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{1}\mspace{14mu}{by}\mspace{11mu}{solving}\mspace{14mu} k_{1}h_{0}} + {k_{0}h_{1}} + k_{0}^{2} + c_{1}} = 0}{\left. c_{2}\leftarrow{f_{3} + {c_{3}u_{20}}} \right.,\left. c_{1}\leftarrow{f_{2} + {h_{2}v_{20}} + {c_{2}u_{20}}} \right.,\left. c_{0}\leftarrow{f_{1} + {h_{1}v_{20}} + {c_{1}u_{20}}} \right.}{\left. \alpha\leftarrow h_{1} \right.,\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right)/\alpha^{2}} \right.}{{{{if}\mspace{14mu}{{Tr}(\gamma)}} = \left. {1\mspace{14mu}{then}\mspace{14mu} k_{1}}\leftarrow k_{1}^{\prime} \right.},\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right)/\alpha^{2}} \right.}{\left. k_{o}\leftarrow{{H(\gamma)}\alpha} \right.,\left. k_{1}^{\prime}\leftarrow{k_{1} + \alpha} \right.}{{{3.\mspace{14mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{0}\mspace{14mu}{by}\mspace{14mu}{checking}\mspace{14mu}{trace}\mspace{14mu}{of}\mspace{14mu}{xh}_{2}} + {x^{2}u_{11}} + 1} = 0}\left. u_{11}\leftarrow\sqrt{{k_{1}h_{1}} + {k_{0}h_{2}} + {k_{1}^{2}u_{20}} + c_{2}} \right.,\left. \gamma\leftarrow{u_{11}/h_{2}^{2}} \right.}{{{if}\mspace{14mu}{{Tr}(\gamma)}} = {1\mspace{14mu}{then}}}{\left. k_{0}\leftarrow k_{0}^{\prime} \right.,\left. u_{11}\leftarrow\sqrt{{k_{1}h_{1}} + {k_{0}h_{2}} + {k_{1}^{2}u_{20}} + c_{2}} \right.}{4.\mspace{14mu}{Compute}\mspace{14mu} U_{1}}\left. u_{10}\leftarrow\sqrt{{k_{0}h_{0}} + {k_{0}^{2}u_{20}} + c_{0}} \right.{{5.\mspace{14mu}{Compute}\mspace{14mu} V_{1}} = {V_{2} + h + {{kU}_{2}\mspace{14mu}{mod}\mspace{14mu} U_{1}}}}\left. w\leftarrow{h_{2} + k_{1}} \right.\left. v_{11}\leftarrow{h_{1} + {k_{1}u_{20}} + k_{0} + {u_{11}w}} \right.\left. v_{10}\leftarrow{v_{20} + h_{0} + {k_{0}u_{20}} + {u_{10}w}} \right.{\left. {6.\mspace{14mu}{U_{1}(x)}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,\left. {V_{1}(x)}\leftarrow{{v_{11}x} + v_{10}} \right.}{{7.\mspace{14mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)}} & \left\lbrack {{Formula}\mspace{14mu} 15} \right\rbrack \end{matrix}$

Processing Example 2 Proposed Method F1

Processing Example 2 (Proposed Method F1) relates to a method of computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+h₁x+h₀, f₄=0.

A close look at Algorithm 5 will reveal that Algorithm 5 contains a large number of multiplication operations by a coefficient h(x) and inversion operations of the coefficient h(x). This means that the complexities of multiplications and inversion operations can be reduced by manipulating the coefficient h(x). It should be noted that according to the document (Non-patent Document 19: T. Lange. Efficient Doubling on Genus Two Curves over Binary Fields, SAC 2004, pre-proceedings, pp. 189-202, 2004.), h₂=1, f₄=0 are used to achieve fast computation. The complexity of HarleyDBL in the case where these parameters are used is 21M+5S+1I.

While the conditions for Processing Example 2 (Proposed Method F1) described here are also set in conformity with those mentioned above, since an irreducible polynomial is assumed for h(x) due to Lemma 1, h(x)=x ² +h ₁ x+h ₀, and Tr(h ₀ /h ₁ ²)=1 are assumed (the necessary and sufficient condition for the quadratic equation ax²+bx+c=0 to be an irreducible polynomial is Tr(ac/b²)=1). The computation method in this case is shown in Algorithm 8 [Algorithm 8] HEC_HLV(h₂=1, f₄=0).

$\begin{matrix} {{{Algorithm}\mspace{14mu} 8\mspace{14mu}{HEC\_ HLV}\mspace{14mu}\left( {{h_{2} = 1},{f_{4} = 0}} \right)}{{{{Input}\text{:}\mspace{14mu} D_{2}} = \left( {U_{2},V_{2}} \right)},{{invu} = {1/u_{21}}},{1/h_{1}^{2}}}{{{{Output}\text{:}\mspace{14mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}},{{invu} = {1/u_{11}}}}{{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},{{\gcd\left( {h,U_{i}} \right)} = 1},{i = 1},2}{{{1.\mspace{14mu}{Solve}\mspace{14mu} k_{1}} + {k_{1}^{2}u_{21}} + 1} = 0}{\left. \alpha\leftarrow{invu} \right.,\left. k_{1}\leftarrow{{H\left( u_{21} \right)}\alpha} \right.,\left. k_{1}^{\prime}\leftarrow{k_{1} + \alpha} \right.}{{{2.\mspace{14mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{1}\mspace{14mu}{by}\mspace{14mu}{solving}\mspace{14mu} k_{1}h_{0}} + {k_{0}h_{1}} + {k_{0}^{2}u_{21}} + c_{1}} = 0}\left. c_{1}\leftarrow{f_{3} + v_{21} + u_{20} + u_{21}^{2}} \right.\left. c_{0}\leftarrow{f_{2} + v_{20} + {v_{21}\left( {h_{1} + v_{21}} \right)} + {u_{21}\left( {u_{20} + c_{1}} \right)}} \right.{\left. w_{0}\leftarrow{u_{21}/h_{1}^{2}} \right.,\left. \alpha\leftarrow{h_{1}\alpha} \right.,\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right)w_{0}} \right.}{{{{if}\mspace{14mu}{{Tr}(\gamma)}} = \left. {1\mspace{14mu}{then}\mspace{14mu} k_{1}}\leftarrow k_{1}^{\prime} \right.},\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right)w_{0}} \right.}{\left. k_{0}\leftarrow{{H(\gamma)}\alpha} \right.,\left. k_{0}^{\prime}\leftarrow{k_{0} + \alpha} \right.}{{{3.\mspace{14mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{0}\mspace{14mu}{by}\mspace{14mu}{checking}\mspace{14mu}{trace}\mspace{14mu}{of}\mspace{14mu} x} + {x^{2}u_{11}} + 1} = 0}{\left. w_{0}\leftarrow k_{1}^{2} \right.,\left. w_{1}\leftarrow{{w_{0}u_{20}} + {k_{1}h_{1}} + u_{21}} \right.}{\left. w_{2}\leftarrow{k_{0} + \sqrt{w_{1} + k_{0}}} \right.,\left. w_{3}\leftarrow{k_{0}^{\prime} + \sqrt{w_{1} + k_{0}^{\prime}}} \right.,\left. w_{4}\leftarrow{w_{2}w_{3}} \right.}{\left. w_{1}\leftarrow{1/\left( {w_{4}k_{1}} \right)} \right.,\left. w_{4}\leftarrow{w_{1}w_{4}} \right.,\left. u_{11}\leftarrow{w_{2}w_{4}} \right.}{{{if}\mspace{14mu}{{Tr}(\gamma)}} = {1\mspace{14mu}{then}}}{\left. k_{0}\leftarrow k_{0}^{\prime} \right.,\left. u_{11}\leftarrow{w_{3}w_{4}} \right.,\left. w_{2}\leftrightarrow w_{3} \right.}\left. {invu}\leftarrow{w_{0}w_{1}w_{3}} \right.{4.\mspace{14mu}{Compute}\mspace{14mu} U_{1}}{\left. w_{1}\leftarrow{k_{0}u_{20}} \right.,\left. w_{5}\leftarrow{k_{1}u_{21}} \right.,\left. w_{6}\leftarrow{\left( {k_{0} + k_{1}} \right)\left( {u_{20} + u_{21}} \right)} \right.}\left. u_{10}\leftarrow{w_{4}\sqrt{{k_{0}\left( {w_{1} + h_{0}} \right)} + c_{0}}} \right.{{5.\mspace{14mu}{Compute}\mspace{14mu} V_{1}} = {V_{2} + h + {{kU}_{2}{mod}\; U_{1}}}}{\left. w_{4}\leftarrow{w_{5} + k_{0} + 1} \right.,\left. w_{5}\leftarrow{w_{1} + w_{5} + w_{6} + v_{21} + h_{1}} \right.}{\left. w_{6}\leftarrow{w_{1} + v_{20} + h_{0}} \right.,\left. w_{7}\leftarrow{{k_{1}u_{11}} + w_{4}} \right.}{\left. w_{1}\leftarrow{w_{7}u_{10}} \right.,\left. w_{3}\leftarrow{\left( {k_{1} + w_{7}} \right)\left( {u_{10} + u_{11}} \right)} \right.}\left. v_{11}\leftarrow{w_{1} + w_{3} + w_{4} + w_{5} + w_{7}} \right.\left. v_{10}\leftarrow{w_{1} + w_{6}} \right.{\left. {6.\mspace{14mu}{U_{1}(x)}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,\left. {V_{1}(x)}\leftarrow{{v_{11}x} + v_{10}} \right.}{{{7.\mspace{14mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)},{invu}}} & \left\lbrack {{Formula}\mspace{14mu} 16} \right\rbrack \end{matrix}$

Further, in order to eliminate the number of inversion operations, a technique called Montgomery trick is used. According to this technique, for example, when it is desired to find the inverse of three finite field elements a, b, and c, first, the product of the three elements is found, and the inverse of this is found as w=1/(a*b*c) or the like. Then, to find the inverse of a, w*b*c is computed. For the inverses of b, c, likewise, w*a*c and w*a*b are computed, respectively.

Typically, the complexity of an inversion is several times higher that of a multiplication (as will be described later, the results of software implementation indicate that complexity of an inversion found is about 8 times higher than that of a multiplication). Accordingly, to find the inverse of three elements, for example, if inversion operation is carried out three times in a straightforward fashion, assuming that I=8M, the resulting complexity is 24M. Conversely, if the above-mentioned Montgomery trick is used, the resulting complexity becomes I+8M=16M, thus enabling faster computation that three inversion operations.

According to Processing Example 2 (Proposed Method F1) being described, the inverse of u₁₁ is found using this Montgomery trick. The inverse of u₁₁ is given as an input for the next halving operation. Accordingly, Algorithm 8 allows computation of [½^(i)]D, and when performing the scalar multiplication of the divisor D, Algorithm 8 can be applied to the right-to-left method, that is, a method in which [½^(i)]D is added. The scalar multiplication using halving operation will be described later. Further, the complexity required at this time is as follows.

(a) If k₁, k₀ are correct values: 24M+2S+1I+3SR+2H+2T

(b) If k₁, k₀′ are correct values: 26M+2S+1I+3SR+2H+2T

(c) If k₁′, k₀ are correct values: 25M+2S+1I+3SR+2H+2T

(d) If k₁′, k₀′ are correct values: 27M+2S+1I+3SR+2H+2T

The averaging of all of the above-mentioned cases (a) to (d) yields 25.5 M+2S+1I+3SR+2H+2T.

The complexity of HarleyDBL was 21M+5S+1I. Here, according to the document [(Non-patent Document 15) E. Knudsen. Elliptic Scalar Multiplication Using Point Halving. ASIACRYPTO '99, LNCS 1716, pp. 135-149, Springer-Verlag, 1999.], it is known that when a finite field is defined by a normal basis, the complexities of S (squaring), SR (square root operation), H (half-trace (operation to find the root of a quadratic equation)), and T (trace (determination as to whether roots exist for a quadratic equation)) can be ignored, and only the complexities of M (multiplication) and I (inversion) need to be taken into account. Therefore, when using a normal basis, Algorithm 8 is slower than that of HarleyDBL by 4.5M.

Further, when a finite field is defined by a polynomial basis, according to the document [(Non-patent Document 16) K. Fong, D. Hankerson, J. Lopez, and A. Menezes. Field inversion operation and point halving revised. Technical Report CORR2003-18, www.cacr.math.uwaterloo.ca/techreports/2003/corr2003-18.pdf18], it is known that in comparison to the complexity of M (multiplication), generally, the complexities of SR (square root operation) and H (half-trace) (operation to find the root of a quadratic equation)) are about SR=H=0.5M. Further, the complexity of T (trace (determination as to whether roots exist for a quadratic equation)) can be ignored. Further, the complexity of S (squaring) is known to be only about several tenths of M (multiplication). However, it is also known that depending on the way in which the polynomial basis is chosen, the complexity of SR may become less than 0.5M. It should be noted that exceptional cases can be computed on the basis of the exceptional cases in Processing Example 1 (proposed Method A1) described above.

Processing Example 3 Proposed Method B1

Processing Example 3 (Proposed Method B1) relates to a method of computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+x+h₀, f₄=0.

As has been described with reference to Processing Example 2 (Proposed Method F1) mentioned above, a close look at the computation algorithm for halving described with reference to Processing Example 1 (Proposed Method A1), that is, Algorithm 5 [Algorithm 5 HEC_HLV] will reveal that Algorithm 5 contains a large number of multiplication operations by a coefficient h(x) and inversion operations of the coefficient h(x). This means that the complexities of multiplications and inversion operations can be reduced by manipulating the coefficient h(x). In the document [J. Pelzl, T. Wollinger, J. Guajardo, and C. Paar. Hyperelliptic curve Cryptosystems: Closing the Performance Gap to Elliptic Curves. Cryptology ePrint Archive, 2003/026, IACR, 2003], there is disclosed an example in which h₂, h₁ε{0, 1}, f₄=0 is used to achieve fast computation.

The complexity of HarleyDBL in the case where these parameters are used is 18M+7S+1I.

While the conditions for Processing Example 3 (Proposed Method B1) are also set in conformity with those mentioned above, since an irreducible polynomial is assumed for h(x) due to Lemma 1 mentioned above, h(x)=x ² +x+h ₀, and Tr(h ₀)=1 are set (the necessary and sufficient condition for the quadratic equation ax²+bx+c=0 to be an irreducible polynomial is Tr(ac/b²)=1).

The computation method in this case is shown below as Algorithm 10 [Algorithm 10 HEC_HLV(h₂=h₁=1, f₄=0).

$\begin{matrix} {{{{Algorithm}\mspace{14mu} 10\mspace{14mu}{HEC\_ HLV}\left( {{h_{2} = {h_{1} = 1}},{f_{4} = 0}} \right)}{{{{Input}\text{:}\mspace{14mu} D_{2}} = \left( {U_{2},V_{2}} \right)},{{invu} = {1/u_{21}}}}{{{{Output}\text{:}\mspace{14mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}},{{invu} = {1/u_{11}}}}{{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},{{\gcd\left( {h,U_{i}} \right)} = 1},{i = 1},2}{{{1.\mspace{14mu}{Solve}\mspace{14mu} k_{1}} + {k_{1}^{2}u_{21}} + 1} = 0}{\left. \alpha\leftarrow{invu} \right.,\left. k_{1}\leftarrow{{H\left( u_{21} \right)}\alpha} \right.,\left. k_{1}^{\prime}\leftarrow{k_{1} + \alpha} \right.}{{{2.\mspace{14mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{1}\mspace{14mu}{by}\mspace{14mu}{solving}\mspace{14mu} k_{1}h_{0}} + k_{0} + {k_{0}^{2}u_{21}} + c_{1}} = 0}\left. c_{1}\leftarrow{f_{3} + v_{21} + u_{20} + u_{21}^{2}} \right.\left. c_{0}\leftarrow{f_{2} + v_{20} + v_{21} + v_{21}^{2} + {u_{21}\left( {u_{20} + c_{1}} \right)}} \right.\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right)u_{21}} \right.{{{{if}\mspace{14mu}{{Tr}(\gamma)}} = \left. {1\mspace{14mu}{then}\mspace{14mu} k_{1}}\leftarrow k_{1}^{\prime} \right.},\left. \gamma\leftarrow{\gamma + h_{0}} \right.}{\left. k_{0}\leftarrow{{H(\gamma)}\alpha} \right.,\left. k_{0}^{\prime}\leftarrow{k_{0} + \alpha} \right.}{{{3.\mspace{14mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{0}\mspace{14mu}{by}\mspace{14mu}{checking}\mspace{14mu}{trace}\mspace{14mu}{of}\mspace{14mu} x} + {x^{2}u_{11}} + 1} = 0}{\left. w_{0}\leftarrow k_{1}^{2} \right.,\left. w_{1}\leftarrow{{w_{0}u_{20}} + k_{1} + u_{21}} \right.}{\left. w_{2}\leftarrow{k_{0} + \sqrt{w_{1} + k_{0}}} \right.,\left. w_{3}\leftarrow{k_{0}^{\prime} + \sqrt{w_{1} + k_{0}^{\prime}}} \right.,\left. w_{4}\leftarrow{w_{2}w_{3}} \right.}{\left. w_{1}\leftarrow{1/\left( {w_{4}k_{1}} \right)} \right.,\left. w_{4}\leftarrow{w_{1}w_{4}} \right.,\left. u_{11}\leftarrow{w_{2}w_{4}} \right.}{{{if}\mspace{14mu}{{Tr}\left( u_{11} \right)}} = {1\mspace{14mu}{then}}}{\left. k_{0}\leftarrow k_{0}^{\prime} \right.,\left. u_{11}\leftarrow{w_{3}w_{4}} \right.,\left. w_{2}\leftrightarrow w_{3} \right.}{4.\mspace{14mu}{Compute}\mspace{14mu} U_{1}}{\left. w_{0}\leftarrow{w_{0}w_{1}} \right.,\left. w_{1}\leftarrow{k_{0}u_{20}} \right.,\left. w_{5}\leftarrow{k_{1}u_{21}} \right.,\left. w_{6}\leftarrow{\left( {k_{0} + k_{1}} \right)\left( {u_{20} + u_{21}} \right)} \right.}\left. u_{10}\leftarrow{w_{4}\sqrt{{k_{0}\left( {w_{1} + h_{0}} \right)} + c_{0}}} \right.{{5.\mspace{14mu}{Compute}\mspace{14mu} V_{1}} = {V_{2} + h + {{kU}_{2}\mspace{14mu}{mod}\mspace{14mu} U_{1}}}}{\left. w_{4}\leftarrow{w_{5} + k_{0} + 1} \right.,\left. w_{5}\leftarrow{w_{1} + w_{5} + w_{6} + v_{21} + 1} \right.,\left. w_{6}\leftarrow{w_{1} + v_{20} + h_{0}} \right.}\left. {invu}\leftarrow{w_{0}w_{3}} \right.\left. w_{0}\leftarrow{w_{2} + w_{4}} \right.,\left. w_{1}\leftarrow{w_{0}u_{10}} \right.,\left. w_{3}\leftarrow{\left( {k_{0} + k_{1}} \right)\left( {u_{10} + u_{11}} \right)} \right.}\left. v_{10}\leftarrow{w_{1} + w_{2} + w_{3} + w_{5}} \right.\left. v_{11}\leftarrow{w_{1} + w_{6}} \right.{\left. {6.\mspace{14mu}{U_{1}(x)}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,\left. {V_{1}(x)}\leftarrow{{v_{11}x} + v_{10}} \right.}{{{7.\mspace{14mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)},{invu}}} & \left\lbrack {{Formula}\mspace{14mu} 17} \right\rbrack \end{matrix}$

Further, in order to eliminate the number of inversion operations, as in Processing Example 2 (Proposed Method F1) described above, the Montgomery trick is used to find the inverse of u₁₁. The inverse of u₁₁ will be given an input for the next halving operation.

The complexity according to this processing example is as follows.

(a) If k₁, k₀ are correct values: 19M+3S+1I+3SR+2H+2T

(b) If k₁, k₀′ are correct values: 20M+3S+1I+3SR+2H+2T

(c) If k₁, k₀′ are correct values: 19M+3S+1I+3SR+2H+2T

(d) If k₁′, k₀′ are correct values: 20M+3S+1I+3SR+2H+2T

The averaging of all of the above-mentioned cases (a) to (d) yields 19.5M+3S+1I+3SR+2H+2T. The complexity of HarleyDBL was 18M+7S+1I. Here, as described above, according to the document [(Non-patent Document 15) E. Knudsen. Elliptic Scalar Multiplication Using Point Halving. ASIACRYPTO '99, LNCS 1716, pp. 135-149, Springer-Verlag, 1999.], it is known that when a finite field is defined by a normal basis, the complexities of S (squaring), SR (square root operation), H (half-trace) (operation to find the root of a quadratic equation)), and T (trace (determination as to whether roots exist for a quadratic equation)) can be ignored, and only the complexities of M (multiplication) and I (inversion) need to be taken into account.

Therefore, when using a normal basis, Algorithm 10 [Algorithm 10] described above is slower than the conventional algorithm [HarleyDBL] by 1.5M. Further, when a finite field is defined by a polynomial basis, according to the document [K. Fong, D. Hankerson, J. Lopez, and A. Menezes. Field inversion operation and point halving revised. Technical Report CORR2003-18, www.cacr.math.uwaterloo.ca/techreports/2003/corr2003-18.pdf18], it is known that generally, the complexities of SR and H are about SR=H=0.5M. Further, the complexity of T can be ignored. Further, it is known that the complexity of S is only about several tenths of M. However, it is also known that depending on the way in which the polynomial basis is chosen, the complexity of SR may become less than 0.5M.

The curve of Algorithm 10 [Algorithm 10] mentioned above is also subject to the constraint h₀=1. Since Algorithm 10 [Algorithm 10] mentioned above involves one multiplication operation of h₀, by setting as h₀=1, the complexity can be reduced by 1M. The complexity found by the averaging of all of the above-mentioned cases (a) to (d) is 18.5M+3S+1I+3SR+2H+2T. On the other hand, the complexity of HarleyDBL is 15M+7S+11. It should be noted that exceptional cases can be computed on the basis of the exceptional cases in Processing Example 1 (proposed Method A1) described above.

Processing Example 4 Proposed Method E1

Processing Example 4 (Proposed Method E1) relates to a method of computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with a parameter h(x)=x.

As in Processing Case 3 (proposed Method B1), in Algorithm 5, by setting as h(x)=x, the complexity of the multiplication and inversion operations of elements on a finite field required for the halving operation of a divisor can be reduced. As a specific example, the algorithm in the case where f(x)=x⁵+f₁x+f₀ is shown below as Algorithm 12 (Algorithm 12).

$\begin{matrix} \begin{matrix} {{Algorithm}\mspace{14mu} 12{HEC\_ HLV}\left( {{{h(x)} = x},{{f(x)} = {x^{5} + {f_{1}x} + f_{0}}}} \right)} \\ {{{{Input}\text{:}\mspace{11mu} D_{2}} = \left( {U_{2},V_{2}} \right)}\;} \\ {{{Output}\text{:}\mspace{11mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}} \\ {\mspace{31mu}\begin{matrix} {{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},} \\ {{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},} \\ {{{\gcd\left( {h,U_{i}} \right)} = 1},{i = 1},2} \end{matrix}} \\ {{{1.\mspace{11mu}{Solve}\mspace{14mu} k_{1}^{2}u_{21}} + 1} = 0} \\ {\mspace{31mu}{\left. w_{0}\leftarrow{1/u_{21}} \right.,\left. k_{1}\leftarrow\sqrt{w_{0}} \right.}} \\ {{{2.\mspace{11mu}{Solve}\mspace{14mu} k_{0}} + {k_{0}^{2}u_{21}} + c_{1}} = 0} \\ {\mspace{25mu}{\left. c_{1}\leftarrow{u_{20} + u_{21}^{2}} \right.,\left. w_{1}\leftarrow{c_{1}u_{21}} \right.}} \\ {\mspace{25mu}\left. c_{0}\leftarrow{v_{21} + v_{21}^{2} + {u_{21}u_{20}} + w_{1}} \right.} \\ {\mspace{25mu}{\left. {invk}_{1}\leftarrow\sqrt{u_{21}} \right.,\left. w_{2}\leftarrow{H\left( w_{1} \right)} \right.,\left. w_{3}\leftarrow{w_{2} + 1} \right.}} \\ {\mspace{25mu}{\left. k_{0}\leftarrow{w_{0}w_{2}} \right.,\left. k_{0}^{\prime}\leftarrow{k_{0} + w_{0}} \right.}} \\ {3.\mspace{11mu}{Compute}\mspace{14mu} U_{1}} \\ {\mspace{25mu}{\left. u_{11}\leftarrow\sqrt{{invk}_{1} + k_{0}} \right.,\left. u_{10}\leftarrow\sqrt{{\left( {k_{0} + c_{1}} \right)u_{20}} + {c_{0}u_{21}}} \right.}} \\ {\mspace{25mu}{{{if}\mspace{14mu}{{Tr}\left( {u_{11}\left( {u_{10} + {invk}_{1} + k_{0}} \right)} \right)}} = {1\mspace{14mu}{then}}}} \\ {\mspace{45mu}\begin{matrix} {\left. k_{0}\leftarrow k_{0}^{\prime} \right.,\left. w_{2}\leftarrow w_{3} \right.,} \\ {\left. u_{11}\leftarrow{u_{11} + k_{1}} \right.,\left. u_{10}\leftarrow{u_{10} + \sqrt{w_{0} + u_{20}}} \right.} \end{matrix}} \\ {{4.\mspace{14mu}{Compute}\mspace{14mu} V_{1}} = {V_{2} + h + {{kU}_{2}\mspace{11mu}{mod}\; U_{1}}}} \\ {\mspace{31mu}\left. w_{1}\leftarrow{{k_{1}\left( {u_{21} + u_{11}} \right)} + k_{0}} \right.} \\ {\mspace{31mu}\left. v_{11}\leftarrow{{k_{1}\left( {u_{20} + u_{10}} \right)} + w_{2} + v_{21} + 1 + {u_{11}w_{1}}} \right.} \\ {\mspace{31mu}\left. v_{10}\leftarrow{{k_{0}u_{20}} + v_{20} + {u_{10}w_{1}}} \right.} \\ {\left. {5.\mspace{14mu}{U_{1}(x)}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,\left. {V_{1}(x)}\leftarrow{{v_{11}x} + v_{10}} \right.} \\ {{6.\mspace{14mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)} \end{matrix} & \left\lbrack {{Formula}\mspace{20mu} 18} \right\rbrack \end{matrix}$

As in Processing Example 3 (Proposed Method B1), the complexity of Algorithm 12 (Algorithm 12) mentioned above is evaluated. Unlike Processing Example 3 (Proposed Method B1), in the case of a hyperelliptic curve of the type where h(x)=x, since k₁ is uniquely determined in step 1, there is only a selection step for k₀ (step 3). The best case with the lowest complexity occurs when Trace in the if sentence in step 3 is 0, and the worst case occurs when Trace is 1. Since the both occur with the same probability, the average complexity is: 11.5M+2S+1I+4.5SR+1H+1T. This complexity is lower than that of Processing Example 3 described above, and thus fast computation is realized. It should be noted that exceptional cases can be computed on the basis of the exceptional cases in Processing Example 1 (proposed Method A1) described above.

Processing Example 5 Proposed Method C1

Processing Example 5 (Proposed Method C1) relates to the method as described below. That is, when computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with random parameters, a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+h₁x+h₀, f₄=0, and a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+x+h₀, f₄=0, two candidates of the halved value arise. In this case, it is necessary to select the one with the correct value from the two candidates. When selecting the correct one, it is necessary to compute the trace, multiplication, and square root of a finite field. Which one of the two candidates is correct depends on the divisor D. Hence, if the divisor D is fixed, information as to which one of the two candidates is correct is retained in a table in advance, and this table is looked up when selecting the correct value, thereby omitting the above-mentioned extra computations.

Which one of k₁, k₁′ (k₀, k₀′) is correct depends on the input divisor D. Accordingly, if D is fixed, for example, when the base point is previously determined as in the case of Phase 1 of ECDH key exchange, ECDSA signature generation or verification, or the like, [½^(i)]D is computed and information as to which of k₁, k₁′ (k₀, k₀′) is correct is recorded in a table in advance.

For example, two tables T₁, T₀ of the same bit size as the order of the base point are prepared, and the binary expression of these tables is represented as: T ₁=(t _(1,r-1), - - - , t _(1,0)), and T ₀=(t _(0,r-1), - - - , t _(0,0)).

When finding [½^(i)]D, if such information that if k₁ is correct, then t_(1,i)=0 or else if k₁ is correct, then t_(1,i)=1; and if k₀ is correct, then t_(1,i)=0 or else if k₀′ is correct, then t_(0,i)=1, is stored in the tables, a bit string only about twice the size of the order of the base point suffices as the table size. By looking up these tables, the complexity of halving can be reduced.

The above-mentioned method as applied to Algorithm 8 [Algorithm 8] HEC_HLV(h₂=1, f₄=0) is represented as Algorithm 9 [Algorithm 9] HEC_HLV(h₁=1, f₄=0, with table-lookup). The complexity of the algorithm is 22M+2SR+1I+2SR+2H.

$\begin{matrix} \begin{matrix} {{Algorithm}\mspace{14mu} 9{HEC\_ HLV}\left( {{h_{2} = 1},{f_{4} = 0},{{table}\text{-}{lookup}}} \right)} \\ {{{{Input}\text{:}\mspace{11mu} D_{2}} = \left( {U_{2},V_{2}} \right)},{{invu} = {1/u_{21}}},{1/h_{1}^{2}},t_{0},t_{1}} \\ {{{{Output}\text{:}\mspace{11mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}},{{invu} = {1/u_{11}}}} \\ {\mspace{34mu}\begin{matrix} {{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},} \\ {{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},} \\ {{{\gcd\left( {h,U_{i}} \right)} = 1},{i = 1},2} \end{matrix}} \\ {{{1.\mspace{11mu}{Solve}\mspace{14mu} k_{1}} + {k_{1}^{2}u_{21}} + 1} = 0} \\ {\mspace{40mu}\begin{matrix} {{{if}\mspace{14mu} t_{1}} = {0\mspace{14mu}{then}}} \\ {\;\left. k_{1}\leftarrow{{H\left( u_{21} \right)}{invu}\mspace{20mu}{else}\mspace{14mu} k_{1}}\leftarrow{{{H\left( u_{21} \right)}{invu}} + {invu}} \right.} \end{matrix}} \\ {{{2.\mspace{11mu}{Solve}\mspace{14mu} k_{1}h_{0}} + {k_{0}h_{1}} + {k_{0}^{2}u_{21}} + c_{1}} = 0} \\ {\mspace{25mu}\left. c_{1}\leftarrow{f_{3} + v_{21} + u_{20} + u_{21}^{2}} \right.} \\ {\mspace{25mu}\left. c_{0}\leftarrow{f_{2} + v_{20} + {v_{21}\left( {h_{1} + v_{21}} \right)} + {u_{21}\left( {u_{20} + c_{1}} \right)}} \right.} \\ {\mspace{25mu}{\left. \alpha\leftarrow{h_{1}{invu}} \right.,\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right){u_{21}/h_{1}^{2}}} \right.}} \\ {\mspace{25mu}{{{if}\mspace{14mu} t_{0}} = \left. {0\mspace{14mu}{then}\mspace{14mu} k_{0}}\leftarrow{{H(\gamma)}\alpha\mspace{14mu}{else}\mspace{14mu} k_{0}}\leftarrow{{{H(\gamma)}\alpha} + \alpha} \right.}} \\ {3.\mspace{11mu}{Compute}\mspace{14mu} U_{1}} \\ {\mspace{25mu}{\left. w_{0}\leftarrow k_{1}^{2} \right.,\left. w_{1}\leftarrow{{w_{0}u_{20}} + {k_{1}h_{1}} + u_{21}} \right.,\left. w_{2}\leftarrow{k_{0} + \sqrt{w_{1} + k_{0}}} \right.}} \\ {\mspace{25mu}{\left. w_{1}\leftarrow{1/\left( {w_{2}k_{1}} \right)} \right.,\left. w_{3}\leftarrow{w_{1}w_{2}} \right.}} \\ {\mspace{25mu}{\left. u_{11}\leftarrow{w_{2}w_{3}} \right.,\left. {invu}\leftarrow{w_{0}w_{1}} \right.}} \\ {\mspace{25mu}{\left. w_{1}\leftarrow{k_{0}u_{20}} \right.,\left. w_{5}\leftarrow{k_{1}u_{21}} \right.,\left. w_{6}\leftarrow{\left( {k_{0} + k_{1}} \right)\left( {u_{20} + u_{21}} \right)} \right.}} \\ {\mspace{25mu}\left. u_{10}\leftarrow{w_{3}\sqrt{{k_{0}\left( {w_{1} + h_{0}} \right)} + c_{0}}} \right.} \\ {{4.\mspace{11mu}{Compute}\mspace{14mu} V_{1}} = {V_{2} + h + {{kU}_{2}\mspace{11mu}{mod}\mspace{11mu} U_{1}}}} \\ {\mspace{25mu}{\left. w_{4}\leftarrow{w_{5} + k_{0} + 1} \right.,\left. w_{5}\leftarrow{w_{1} + w_{5} + w_{6} + v_{21} + h_{1}} \right.}} \\ {\mspace{25mu}{\left. w_{6}\leftarrow{w_{1} + v_{20} + h_{0}} \right.,\left. w_{7}\leftarrow{{k_{1}u_{11}} + w_{4}} \right.}} \\ {\mspace{25mu}{\left. w_{3}\leftarrow{\left( {k_{1} + w_{7}} \right)\left( {u_{10} + u_{11}} \right)} \right.,\left. w_{1}\leftarrow{w_{7}u_{10}} \right.}} \\ {\mspace{25mu}\left. v_{11}\leftarrow{w_{1} + w_{3} + w_{4} + w_{5} + w_{7}} \right.} \\ {\mspace{25mu}\left. v_{10}\leftarrow{w_{1} + w_{6}} \right.} \\ {\left. {6.\mspace{11mu}{U_{1}(x)}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,\left. {V_{1}(x)}\leftarrow{{v_{11}x} + v_{10}} \right.} \\ {{{7.\mspace{11mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)},{invu}} \end{matrix} & \left\lbrack {{Formula}\mspace{20mu} 19} \right\rbrack \end{matrix}$

Specifically, the above-mentioned method as applied to Algorithm 10 [Algorithm 10 HEC_HLV(h₂=h₁=1, f₄=0] described above is represented as Algorithm 11 [Algorithm 11 HEC_HLV(h₂=h₁=1, f₄=0, with table-lookup) below.

$\begin{matrix} \begin{matrix} {{Algorithm}\mspace{14mu} 11{HEC\_ HLV}\begin{pmatrix} {{h_{2} = {h_{1} = 1}},} \\ {{f_{4} = 0},{{with}\mspace{14mu}{table}\text{-}{lookup}}} \end{pmatrix}} \\ {{{{Input}\text{:}\mspace{11mu} D_{2}} = \left( {U_{2},V_{2}} \right)},{{invu} = {1/u_{21}}},t_{0},t_{1}} \\ {{{{Output}\text{:}\mspace{11mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}},{{invu} = {1/u_{11}}}} \\ {\mspace{31mu}{{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},{{\gcd\left( {h,U_{1}} \right)} = 1}}} \\ {{{1.\mspace{11mu}{Solve}\mspace{14mu} k_{1}} + {k_{1}^{2}u_{21}} + 1} = {0\mspace{20mu}{and}\mspace{14mu}{select}\mspace{14mu}{correct}\mspace{14mu} k_{1}\mspace{11mu}{via}\mspace{14mu} t_{1}}} \\ {\mspace{31mu}{\left. \alpha\leftarrow{invu} \right.,\left. k_{1}\leftarrow{{H\left( u_{21} \right)}\alpha} \right.}} \\ {\mspace{31mu}{{{if}\mspace{14mu} t_{1}} = \left. {0\mspace{14mu}{then}\mspace{20mu} k_{1}}\leftarrow{k_{1} + \alpha} \right.}} \\ {{{{2.\mspace{14mu}{Solving}\mspace{14mu} k_{1}h_{0}} + k_{0} + {k_{0}^{2}u_{21}} + c_{1}} = 0}\mspace{14mu}} \\ {\mspace{20mu}\begin{matrix} {{and}\mspace{14mu}{select}\mspace{14mu}{correct}\mspace{14mu} k_{0}\mspace{14mu}{via}\mspace{14mu} t_{0}} \\ \left. c_{1}\leftarrow{f_{3} + v_{21} + u_{20} + u_{21}^{2}} \right. \end{matrix}\;} \\ {\mspace{25mu}\left. c_{0}\leftarrow{f_{2} + v_{20} + v_{21} + v_{21}^{2} + {u_{21}\left( {u_{20} + c_{1}} \right)}} \right.} \\ {\mspace{25mu}\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right)u_{21}} \right.} \\ {\mspace{25mu}\left. k_{0}\leftarrow{{H(\gamma)}\alpha} \right.} \\ {\mspace{25mu}{{{if}\mspace{14mu} t_{1}} = \left. {0\mspace{14mu}{then}\mspace{14mu} k_{0}}\leftarrow{k_{0} + \alpha} \right.}} \\ {3.\mspace{11mu}{Compute}\mspace{14mu} U_{1}} \\ {\mspace{25mu}{\left. w_{0}\leftarrow k_{1}^{2} \right.,\left. w_{1}\leftarrow{{w_{0}u_{20}} + k_{1} + u_{21}} \right.}} \\ {\mspace{25mu}\left. w_{2}\leftarrow{k_{0} + \sqrt{w_{1} + k_{0}}} \right.} \\ {\mspace{25mu}{\left. w_{1}\leftarrow{1/\left( {w_{2}k_{1}} \right)} \right.,\left. w_{4}\leftarrow{w_{1}w_{2}} \right.}} \\ {\mspace{25mu}\left. u_{11}\leftarrow{w_{2}w_{4}} \right.} \\ {\mspace{25mu}\begin{matrix} {\left. w_{0}\leftarrow{w_{0}w_{1}} \right.,\left. w_{1}\leftarrow{k_{0}u_{20}} \right.,} \\ {\left. w_{5}\leftarrow{k_{1}u_{21}} \right.,\left. w_{6}\leftarrow{\left( {k_{0} + k_{1}} \right)\left( {u_{20} + u_{21}} \right)} \right.} \end{matrix}} \\ {\mspace{25mu}\left. u_{10}\leftarrow{w_{4}\sqrt{{k_{0}\left( {w_{1} + h_{0}} \right)} + c_{0}}} \right.} \\ {{5.\mspace{11mu}{Compute}\mspace{14mu} V_{1}} = {V_{2} + h + {{kU}_{2}\mspace{11mu}{mod}\mspace{11mu} U_{1}}}} \\ {\mspace{20mu}\begin{matrix} {\;{\left. w_{4}\leftarrow{w_{5} + k_{0} + 1} \right.,}} \\ {\left. w_{5}\leftarrow{w_{1} + w_{5} + w_{6} + v_{21} + 1} \right.,\left. w_{6}\leftarrow{w_{1} + v_{20} + h_{0}} \right.} \end{matrix}} \\ {\mspace{25mu}\left. {invu}\leftarrow w_{0} \right.} \\ {\mspace{25mu}{\left. w_{0}\leftarrow{w_{2} + w_{4}} \right.,\left. w_{1}\leftarrow{w_{0}u_{10}} \right.,\left. w_{3}\leftarrow{\left( {k_{0} + k_{1}} \right)\left( {u_{10} + u_{11}} \right)} \right.}} \\ {\mspace{25mu}\left. v_{10}\leftarrow{w_{1} + w_{2} + w_{3} + w_{5}} \right.} \\ {\mspace{25mu}\left. v_{11}\leftarrow{w_{1} + w_{6}} \right.} \\ {\left. {6.\mspace{11mu}{U_{1}(x)}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,\left. {V_{1}(x)}\leftarrow{{v_{11}x} + v_{10}} \right.} \\ {{{7.\mspace{11mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)},{{invu} = {1/u_{11}}}} \end{matrix} & \left\lbrack {{Formula}\mspace{20mu} 20} \right\rbrack \end{matrix}$

The complexity of Algorithm 11 [Algorithm 11 HEC_HLV(h₂=h₁=1, f₄=0, with table-lookup) is 18M+3S+1I+2SR+2H, and further, by setting h₀=1, the complexity can be reduced by 1M. The complexity in this case becomes 17M+3S+1I+2SR+2H.

Further, the above-mentioned method as applied to Algorithm 12 [HEC_HLV(h(x)=x, f(X)=x⁵+f₁x+f₀] described above is represented as Algorithm 13 [HEC_HLV(h(x)=x, f(X)=x⁵+f₁x+f₀, with table-lookup) below.

$\begin{matrix} \begin{matrix} {{Algorithm}\mspace{14mu} 13{HEC\_ HLV}\begin{pmatrix} {{{h(x)} = x},{{f(x)} = {x^{5} + {f_{1}x} + f_{0}}},} \\ {{with}\mspace{14mu}{table}\text{-}{lookup}} \end{pmatrix}} \\ {{{{Input}\text{:}\mspace{11mu} D_{2}} = \left( {U_{2},V_{2}} \right)},t_{0}} \\ {{{Output}\text{:}\mspace{11mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}} \\ {\mspace{31mu}{{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},{{\gcd\left( {h,U_{1}} \right)} = 1}}} \\ {{{1.\mspace{11mu}{Solve}\mspace{14mu} k_{1}^{2}u_{21}} + 1} = 0} \\ {\mspace{25mu}{\left. w_{0}\leftarrow{1/u_{21}} \right.,\left. k_{1}\leftarrow\sqrt{w_{0}} \right.}} \\ {{{2.\mspace{14mu}{Solve}\mspace{14mu} k_{0}} + {k_{0}^{2}u_{21}} + c_{1}} = 0} \\ {\mspace{31mu}{\left. c_{1}\leftarrow{u_{20} + u_{21}^{2}} \right.,\left. w_{1}\leftarrow{c_{1}u_{21}} \right.}} \\ {\mspace{31mu}\left. c_{0}\leftarrow{v_{21} + v_{21}^{2} + {u_{21}u_{20}} + w_{1}} \right.} \\ {\mspace{31mu}{\left. {invk}_{1}\leftarrow\sqrt{u_{21}} \right.,\left. w_{2}\leftarrow{H\left( w_{1} \right)} \right.}} \\ {\mspace{31mu}{{{if}\mspace{14mu} t_{0}} = {0\mspace{14mu}{then}}}} \\ {\mspace{85mu}\left. k_{0}\leftarrow{w_{0}w_{2}} \right.} \\ {\mspace{59mu}{else}} \\ {\mspace{85mu}{\left. k_{0}\leftarrow{{w_{0}w_{2}} + w_{0}} \right.,\left. w_{2}\leftarrow{w_{2} + 1} \right.}} \\ {3.\mspace{14mu}{Compute}\mspace{14mu} U_{1}} \\ {\mspace{31mu}{\left. u_{11}\leftarrow\sqrt{{invk}_{1} + k_{0}} \right.,\left. u_{10}\leftarrow\sqrt{{\left( {k_{0} + c_{1}} \right)u_{20}} + {c_{0}u_{21}}} \right.}} \\ {{4.\mspace{11mu}{Compute}\mspace{14mu} V_{1}} = {V_{2} + h + {{kU}_{2}\mspace{11mu}{mod}\mspace{11mu} U_{1}}}} \\ {\mspace{25mu}\left. w_{1}\leftarrow{{k_{1}\left( {u_{21} + u_{11}} \right)} + k_{0}} \right.} \\ {\mspace{25mu}\left. v_{11}\leftarrow{{k_{1}\left( {u_{20} + u_{10}} \right)} + w_{2} + v_{21} + 1 + {u_{11}w_{1}}} \right.} \\ {\mspace{25mu}\left. v_{10}\leftarrow{{k_{0}u_{20}} + v_{20} + {u_{10}w_{1}}} \right.} \\ {\left. {5.\mspace{11mu}{U_{1}(x)}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,\left. {V_{1}(x)}\leftarrow{{v_{11}x} + v_{10}} \right.} \\ {{6.\mspace{11mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)} \end{matrix} & \left\lbrack {{Formula}\mspace{20mu} 21} \right\rbrack \end{matrix}$

The complexity of the above-mentioned algorithm is 9.5M+3S+1I+3.5SR+1H, and thus faster computation can be realized.

Processing Example 6 Proposed Method D1

Processing Example 6 (Proposed Method D1) relates to a method of computing the scalar multiplication of a divisor by using the method of computing the halving of a divisor as set forth in each of Processing Examples 1 to 5.

A method of computing the scalar multiplication using halving of a rational point on an elliptic curve is described in the document [E. Knudsen. Elliptic Scalar Multiplication Using Point Halving. ASIACRYPTO '99, LNCS 1716, pp. 135-149, Springer-Verlag, 1999]. The computation method of scalar multiplication using halving of a divisor on a hyperelliptic curve is executed on the basis of the scalar multiplication disclosed in this document. It should be noted, however, that the right-to-left method, in which [½^(i)]D is computed and added, is used. This algorithm is represented below as Algorithm 14 [Scalar Multiplication].

$\begin{matrix} \begin{matrix} {{Algorithm}\mspace{14mu} 14\mspace{14mu}{Scalar}\mspace{14mu}{Multiplication}} \\ {\begin{matrix} {{{{Input}\text{:}\mspace{11mu} D} \in {{{J\left( {??}_{2^{n}} \right)}\mspace{14mu}{such}\mspace{14mu}{that}\mspace{14mu} 2D} \neq O}},{d \in {\mathbb{Z}}},} \\ {{r\text{:}\mspace{11mu}{order}\mspace{14mu}{of}\mspace{14mu} D},{m = \left\lfloor {\log_{2}r} \right\rfloor}} \end{matrix}\mspace{11mu}} \\ {{Output}\text{:}\mspace{11mu}{scalar}\mspace{14mu}{multiplication}\mspace{14mu}{dD}} \\ {\left. {1.\mspace{11mu}{\sum\limits_{i = 0}^{m}{{\hat{d}}_{i}2^{i}}}}\leftarrow{2^{m}d\mspace{11mu}{mod}\mspace{11mu} r} \right.,{{\hat{d}}_{i} \in \left\{ {0,1} \right\}}} \\ {\left. {2.\mspace{11mu}{\sum\limits_{i = 0}^{m}{d_{i}/2^{i}}}}\leftarrow{\sum\limits_{i = 0}^{m}{{\hat{d}}_{i}2^{i - m}}} \right.,{d \in \left\{ {0,1} \right\}}} \\ {\left. {3.\mspace{11mu} Q}\leftarrow O \right.,\left. R\leftarrow D \right.,\left. {invu}\leftarrow{1/u_{1}} \right.} \\ {4.\mspace{11mu}{for}\mspace{14mu} i\mspace{14mu}{from}\mspace{14mu} 0\mspace{14mu}{to}\mspace{20mu} m\mspace{14mu}{do}\text{:}} \\ {\mspace{31mu}{{{if}\mspace{14mu} d_{i}} = \left. {1\mspace{14mu}{then}\mspace{14mu} Q}\leftarrow{Q + R} \right.}} \\ {\mspace{31mu}\left. \left( {R,{invu}} \right)\leftarrow{{HEC\_ HLV}\left( {R,{invu}} \right)} \right.} \\ {5.\mspace{11mu}{return}\mspace{14mu} Q} \end{matrix} & \left\lbrack {{Formula}\mspace{20mu} 22} \right\rbrack \end{matrix}$

HEC_HLV appearing in step 4 of Algorithm 14 [Algorithm 14 Scalar Multiplication] mentioned above may be HEC_HLV of Algorithm 5 [Algorithm 5] described above using a random curve, HEC_HLV with constraints h₂=1, f₄=0 provided to the curve parameters of Algorithm 8 [Algorithm 8], HEC_HLV with the table-lookup method applied to the curve parameters of Algorithm 8 [Algorithm 8], HEC_HLV with constraints h₂=h₁=1, f₄=0 provided to the curve parameters of Algorithm 10 [Algorithm 10], HEC_HLV with constraints h₂=h₁=h₀=1, f₄=0 provided to the curve parameters of Algorithm 10 [Algorithm 10], or HEC_HLV with the table-lookup method applied to the curve parameters of Algorithm 10 [Algorithm 10]. Further, the above-mentioned HEC_HLV may be HEC_HLV of the curve parameters of Algorithm 12 [Algorithm 12], or HEC_HLV with the table-lookup method applied to Algorithm 12 [Algorithm 12].

[Verification of Increased Computation Speed]

Next, the complexity of the computation applied to each of Processing Examples 1 to 6 described above is found, and verification is made as to an increase in computation speed.

In the case of HEC_HLV(h₂=1, f₄=0), the required complexity is, on average, 25.5M+2S+1I+3SR+2H+2T.

First, a case where a finite field is defined by a normal basis is considered. As described above, when using a normal basis, only the complexity of M and I may be taken into account. According to the document [A. Menezes. Elliptic Curve Public Key Cryptosystems. Kluwer Academic Publishers, 1993.], assuming that finite fields are Fq, q=2^(n), one inversion operation is equivalent to the number of multiplication operations computed by the following expression, that is: └ log₂(n−1)┘+w(n−1)−1  [Formula 23]

In this case, w(n−1) denotes the number of 1's in the binary representation of n−1. For example, if n=83, 89, 113, then I=8M, and if n=103, then I=9M.

Here, assuming that I=8M, the complexity of HEC_HLV(h ₂=1,f ₄=0) is represented as 25.5M+1I=33.5M.

On the other hand, in the case of HarleyDBL, its complexity is represented as 21M+1I=29M, so HarleyDBL is about 13% faster than HEC_HLV. Further, when the table-lookup method is used, the complexity becomes 22M+1I=30M, so HarleyDBL is about 3% faster than HEC_HLV.

Further, in the case of HEC_HLV(h₂=h₁=1, f₄=0), the complexity is, on average, 19.5M+3S+1I+3SR+2H+2T. In this case, 19.5M+1I=27.5M.

On the other hand, in the case of HarleyDBL, the complexity is represented as 18M+1I=26M, so HarleyDBL is about 5% faster than HEC_HLV. Further, when the table-lookup method is used, the complexity becomes 18M+1I=26M, so HarleyDBL and HEC_HLV are equal in complexity.

Further, the complexity of HEC_HLV(h₂=h₁=h₀=1, f₄=0) is, on average, 18.5M+3S+1I+3SR+2H+2T. In this case, 18.5M+1I=26.5M.

On the other hand, in the case of HarleyDBL, its complexity is 15M+1I=23M, so HarleyDBL is about 13% faster than HEC_HLV. Further, when the table-lookup method is used, the complexity becomes 17M+1I=25M, so HarleyDBL is about 8% faster than HEC_HLV.

Further, speed comparison was carried out for the case of a polynomial basis through software implementation.

The software implementation was carried out under the environment as indicated below:

CPU: PentiumII 300 MHx

OS: RedHat7.3

Compiler: gcc2.96.

The operations of M (multiplication) and S (squaring), I (inversion), SR (square root operation) and T (trace (determination as to whether roots exist for a quadratic equation)), and H (half-trace) (operation to find the root of a quadratic equation)) were carried out in the manner as disclosed in the following documents: [D. Hankerson, J. Hernandez, and A. Menezes. Software Implementation of Elliptic Curve Cryptography over Binary Fields. CHES 2000, LNCS 1965, pp. 1-24, 2000. Algorithm 4.6, 4.7]; [S. Shantz. From Euclid's GCD to Montgomery Multiplication to the Great Divide. TR-2001-95, Sun Microsystems, Inc., 2001.]; [K. Fong, D. Hankerson, J. Lopez, and A. Menezes. Field inversion operation and point halving revised. Technical Report CORR2003-18, www.cacr.math.uwaterloo.ca/techreports/2003/corr2003-18.pdf]; and [K. Fong, D. Hankerson, J. Lopez, and A. Menezes. Field inversion operation and point halving revised. Technical Report CORR2003-18, www.cacr.math.uwaterloo.ca/techreports/2003/corr2003-18.pdf Algorithm 4.7], respectively.

M, S, I, SR, H, T were implemented with respect to three finite fields of n=83, 89, 113, and the ratios to M were found. In this case, the following irreducible polynomials were used:

in the case of n=3, z ⁸³ +z ⁷ +z ⁴ +z ²+1=0;

in the case of n=9, z ⁸⁹ +z ³⁸+1=0; and

in the case of n=113, z ¹¹³ +z ⁹+1=0.

The complexities in the respective cases were as follows.

n=83: S/M=0.12, I/M=7.96, SR/M=0.57, H/M=0.58

n=89: S/M=0.05, I/M=8.74, SR/M=0.14, H/M=0.61

n=113: S/M=0.06, I/M=8.56, SR/M=0.10, H/M=0.50

Applying these to the complexity of HarleyDBL, 21M+5S+1I, yields the following.

n=83: HarleyDBL 29.56M

n=89: HarleyDBL 29.99M

n=113: HarleyDBL 29.86M

Applying these to the complexity of HEC_HLV(h₂=1, f₄=0), 25.5M+2S+1I+3SR+2H+2T, yields the following.

n=83: HEC_HLV(h₂=1, f₄=0) 36.57M

n=89: HEC_HLV(h₂=1, f₄=0) 35.98M

n=113: HEC_HLV(h₂=1, f₄=0) 35.48M

In this case, when n=83, 89, 113, HarleyDBL is 20%, 17%, 16% faster than HEC_HLV, respectively.

Further, applying these to the complexity of HEC_HLV(h₂=1, f₄=0) to which the table-lookup method is applied, 22M+2S+1I+2SR+2H, yields the following.

n=83: HEC_HLV(h₂=1, f₄=0 with table-lookup) 32.5M

n=89: HEC_HLV(h₂=1, f₄=0 with table-lookup) 32.34M

n=113: HEC_HLV(h₂=1, f₄=0 with table-lookup) 31.88M

In this case, when n=83, 89, 113, HarleyDBL is 9%, 7%, 6% faster than HEC_HLV, respectively.

Further, in the case of h₂=h₁=1, f₄=0, applying these to the complexity of HarleyDBL, 18M+7S+1I, yields the following.

n=83: HarleyDBL 27.4M

n=89: HarleyDBL 27.09M

n=113: HarleyDBL 26.98M

Next, applying these to the complexity of HEC_HLV(h₂=h₁=1, f₄=0), 19.5M+3S+1I+3SR+2H+2T, yields the following.

n=83: HEC_HLV(h₂=h₁=1, f₄=0) 30.69M

n=89: HEC_HLV(h₂=h₁=1, f₄=0) 30.03M

n=113: HEC_HLV(h₂=h₁=1, f₄=0) 29.54M

In this case, when n=83, 89, 113, HarleyDBL is 13%, 12%, 9% faster than HEC_HLV, respectively.

Further, applying these to the complexity of HEC_HLV(h₂=h₁=1, f₄=0) to which the table-lookup method is applied, 18M+3S+1I+2SR+2H, yields the following.

n=83: HEC_HLV(h₂=h₁=1, f₄=0 with table-lookup) 28.62M

n=89: HEC_HLV(h₂=h₁=1, f₄=0 with table-lookup) 28.39M

n=113: HEC_HLV(h₂=h₁=1, f₄=0 with table-lookup) 27.94M

In this case, when n=83, 89, 113, HarleyDBL is 4%, 5%, 3% faster than HEC_HLV, respectively.

Further, in the case of h₂=h₁=h₀=1, f₄=0, applying these to the complexity of HarleyDBL, 15M+7S+1I, yields the following.

n=83: HarleyDBL 23.8M

n=89: HarleyDBL 24.09M

n=113: HarleyDBL 23.98M

Next, applying these to the complexity of HEC_HLV(h₂=h₁ h₀=1, f₄=0), 18.5M+3S+1I+3SR+2H+2T, yields the following.

n=83: HEC_HLV(h₂=h₁=h₀=1, f₄=0) 29.69M

n=89: HEC_HLV(h₂=h₁=h₀=1, f₄=0) 29.03M

n=113: HEC_HLV(h₂=h₁=h₀=1, f₄=0) 28.54M

In this case, when n=83, 89, 113, HarleyDBL is 20%, 17%, 16% faster than HEC_HLV, respectively.

Further, applying these to the complexity of HEC_HLV(h₂=h₁=h₀=1, f₄=0) to which the table-lookup method is applied, 17M+3S+1I+2SR+2H, yields the following.

n=83: HEC_HLV(h₂=h₁=h₀=1, f₄=0 with table-lookup) 27.62M

n=89: HEC_HLV(h₂=h₁=h₀=1, f₄=0 with table-lookup) 27.39M

n=113: HEC_HLV(h₂=h₁=h₀=1, f₄=0 with table-lookup) 26.94M

In this case, when n=83, 89, 113, HarleyDBL is 14%, 12%, 11% faster than HEC_HLV, respectively.

Next, a comparison with HarleyDBL is made with respect to each of Algorithm 12 [Algorithm 12] described above, that is, [Algorithm 12 HEC_HLV(h(x)=x, f(X)=x⁵+f₁x+f₀)], and Algorithm 13 [Algorithm 13] to which the table-lookup method is applied, that is, [Algorithm 13 HEC_HLV(h(x)=x, f(X)=x⁵+f₁x+f₀, with table-lookup)].

The complexity of [Algorithm 12 HEC_HLV(h(x)=x, f(X)=x⁵+f₁x+f₀)] is 11.5M+2S+1I+4.5SR+1H+1T, and the complexity of [Algorithm 13 HEC_HLV(h(x)=x, f(X)=x⁵+f₁x+f₀, with table-lookup)] is 9.5M+2S+1I+3.5SR+1H. According to the document [(Non-patent Document 19: T. Lange. Efficient Doubling on Genus Two Curves over Binary Fields, SAC 2004, pre-proceedings, pp. 189-202, 2004.)], the complexity of HarleyDBL is 6M+5S+1I. As described above, when a finite field is defined by a normal basis, the complexities of S (squaring), SR (square root operation), H (half-trace (operation to find the root of a quadratic equation)), and T (trace (determination as to whether roots exist for a quadratic equation)) can be ignored, and only the complexities of M (multiplication) and I (inversion) need to be taken into account.

Therefore, the complexity of Algorithm 12 [Algorithm 12 HEC_HLV(h(x)=x, f(X)=x⁵+f₁x+f₀)] becomes 11.5M+2S+1I+4.5SR+1H+1T, =11.5M+1I. Further, the complexity of Algorithm 13 [Algorithm 13 HEC_HLV(h(x)=x, f(X)=x⁵+f₁x+f₀, with table-lookup)] becomes 9.5M+2S+1I+3.5SR+1H =9.5M+1I.

The complexity of HarleyDBL becomes 6M+1I. Therefore, HarleyDBL is faster than HEC_HLV.

As described above, in the case of curve parameters h₂=h₁=1, f₄=0, the complexity in the case where the table-look-up method is applied to HEC_HLV, 18M+3S+1I+2SR+2H, is substantially equal to the complexity of HarleyDBL, 18M+7S+1I, and represents the fastest algorithm upon comparison between HEC_HLV and HarleyDBL under equivalent conditions.

Next, the complexity of scalar multiplication using Algorithm 14 [Algorithm 14 scalar multiplication] is considered. The complexity of scalar multiplication is considered with respect to a method in which h₂=h₁=1, f₄=0 allowing the fastest computation in comparison to HarleyDBL under equivalent conditions are used as the curve parameters, and the table look-up method is used for HEC_HVL.

Since the ratio of steps 1, 2 to the entire scalar multiplication process in Algorithm 14 is very small, the complexity thereof is ignored. Here, the complexity is considered for the cases of n=83, 89, 113 for both a normal basis and a polynomial basis. Further, the order of the base point is assumed to be 165 bits, 177 bits, 225 bits with respect to n=83, 89, 113, respectively. Further, in the repeating portion of step 4, the repetition is made for the number of bits of the order of the base point. Divisor addition is carried out in the manner as disclosed in the document [T. Lange, Efficient arithmetic on genus 2 hyperelliptic curves over finite fields via explicit formulae. Cryptology ePrint Archive, 2002/121, IACR, 2002]. It should be noted that the curve parameters are h₂=h₁=1, f₄=0. The complexity required for the divisor addition in this case is 21M+3S+1I. It is assumed that binary expression of the scalar value results in the appearance of 0, 1 at equivalent ratios. The complexity is computed as follows: ((the complexity of addition)/2+(the complexity of halving or doubling))×(the number of bits of the order of the base point). First, the case of a normal basis will be considered. It is assumed that I=8M.

In the case of h₂=h₁=1, f₄=0,

n=83: addition•doubling: 6682.5M

n=89: addition•doubling: 7168.5M

n=113: addition•doubling: 9112.5M

In the case of h₂=h₁=1, f₄=0,

n=83: addition•halving: 6930M

n=89: addition•halving: 7434M

n=113: addition•halving: 9450M

In the case of h₂=h₁=1, f₄=0+table loop-up method (the complexity is equal to that in the case of addition•doubling)

n=83: addition•doubling: 6682.5M

n=89: addition•doubling: 7168.5M

n=113: addition•doubling: 9112.5M

Next, the case of a polynomial is considered.

In the case of h₂=h₁=1, f₄=0,

n=83: addition•doubling: 6913.5M

n=89: addition•doubling: 7361.3M

n=113: addition•doubling: 9333M

In the case of h₂=h₁=1, f₄=0,

n=83: addition•halving: 7456.35M

n=89: addition•halving: 7881.8M

n=113: addition•halving: 9909M

In the case of h₂=h₁=1, f₄=0+table loop-up method (the complexity is equal to that in the case of addition•doubling)

n=83: addition•doubling: 7114.8M

n=89: addition•doubling: 7591.53M

n=113: addition•doubling: 9540M

As has been described above, according to the processing examples of the present invention described above, halving on elliptic curve cryptography is extended to hyperelliptic curve cryptography to thereby realize fast computation. In the case of cryptographic computation employing computations on a divisor on a hyperelliptic curve, an arithmetic computation that puts a large load on the processing is the scalar multiplication of a divisor. In this regard, the processing according to the present invention as described above enables the scalar multiplication to be computed at a speed equivalent to that of the related art. As a result, even when using halving, hyperelliptic curve cryptography can be processed at a speed equivalent to that of the related art.

Next, Processing Examples 7 to 11 representing improvements over the processing examples described above, that is,

(Processing Example 1: Proposed Method A1)

(Processing Example 2: Proposed Method F1)

(Processing Example 3: Proposed Method B1)

(Processing Example 5: Proposed Method C1)

(Processing Example 6: Proposed Method D1), will be described. Specifically, Processing Examples 7 to 11 refer to the following methods.

(Processing Example 7: Proposed Method A2): A technique aimed at a further increase in the operation speed of the processing example mentioned above (Processing Example 1: Proposed Method A1), which includes computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with random parameters.

(Processing Example 8: Proposed Method F2): A technique aimed at a further increase in the operation speed of Processing Example mentioned above (Processing Example 2: Proposed Method F1), which includes computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+h₁x+h₀, f₄=0

(Processing Example 9: Proposed Method B2): A technique aimed at a further increase in the operation speed of Processing Example mentioned above (Processing Example 3: Proposed Method B1), which includes computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+x+h₀, f₄=0

(Processing Example 10: Proposed Method C2): A technique aimed at a further increase in the operation speed of Processing Example mentioned above (Processing Example 5: Proposed Method C1). That is, when computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with random parameters, a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+h₁x+h₀, f₄=0, and a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+x+h₀, f₄=0, two candidates of the halved value arise. In this case, it is necessary to select the one with the correct value from the two candidates. When selecting the correct one, it is necessary to compute the trace, multiplication, and square root of a finite field. Which one of the two candidates is correct depends on the divisor D. Hence, if the divisor D is fixed, information as to which one of the two candidates is correct is retained in a table in advance, and this table is looked up when selecting the correct value, thereby omitting the above-mentioned extra computations.

(Processing Example 11: Proposed Method D2): A technique aimed at a further increase in the operation speed of Processing Example mentioned above (Processing Example 6: Proposed Method D1), which includes computing the scalar multiplication of a divisor by using the method of computing the halving of a divisor as set forth in each of Processing Examples 7 to 10.

The respective processing examples mentioned above will be sequentially described below in detail.

Processing Example 7 Proposed Method A2

Processing Example 7 (Proposed Method A2) relates to a technique aimed at a further increase in the operation speed of Processing Example mentioned above (Processing Example 1: Proposed Method A1), which includes computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with random parameters.

In the following processing examples as well, it is assumed that the order of the divisors used in the processing below is r. That is, the divisors have no ramification point. It is assumed that the input divisors are as follows: D ₂=(U ₂ ,V ₂) U ₂ =u ₂₂ x ² +u ₂₁ x+u ₂₀; and V ₂ =v ₂₁ x+v ₂₀, where:

u₂₂=1 if the weight of the divisor D₂ is 2; and

u₂₂=1, U₂₁=1, and v₂₁=0 if the weight of the divisor D₂ is 1.

Since no ramification point is included, as the halving operation, it suffices to consider four inversion operations, ExHarDBL^(1+1→2), ExHarDBL^(2+2→1), ExHarDBL^(2+2→2), and HarleyDBL. Those other than HarleyDBL represent exceptional cases.

Here, ExHarDBL^(2+2→1) represents a computation in the case where the weight of the input divisor is 2, and the weight of the output divisor is 1. Further, ExHarDBL^(2+2→2) represents a computation in the case where the weight of the input divisor is 2, and the coefficient of the first order term of U₂ satisfies u₂₁=0 and the weight of the output divisor is 2. However, while ExHarDBL^(2+2→2) can be computed by HarleyDBL, since halving as the inversion operation thereof becomes an exceptional case, ExHarDBL^(2+2→2) is herein treated as an exceptional case.

The halving operations corresponding to ExHarDBL^(1+1→2), ExHarDBL^(2+2→1), ExHarDBL^(2+2→2), and HarleyDBL mentioned above are defined as ExHEC_HLV^(2→1+1), ExHEC_HLV^(1→2+2), ExHEC_HLV^(2→2+2), and HEC_HLV, respectively.

When carrying out the halving operation of a divisor, first, as described above with reference to (Processing Example 1: Proposed Method A), a case-by-case differentiation is made depending on the input divisor as shown in FIG. 2. If the weight of the input divisor is 2, and the coefficient of the first order term of U₂ satisfies u₂₁≠0, the computation is performed by HEC_HLV. Further, if the weight of the input divisor is 2, and the coefficient of the term of U₂ satisfies u₂₁=0, the computation is performed by ExHEC_HLV^(2→2+2) or ExHEC_HLV^(2→1+1). Further, if the weight of the input divisor is 1, the computation is performed by ExHEC_HLV^(1→2+2). The algorithm of HEC_HLV is as descried above with reference to FIG. 3.

The halving operation of a divisor is realized by the reverse operation of the algorithm for performing the doubling operation of a divisor, that is, the [Algorithm 1 Harley DBL] below.

$\begin{matrix} \begin{matrix} {{Algorithm}\mspace{14mu} 1\mspace{14mu}{HarleyDBL}} \\ {{{Input}\text{:}\mspace{11mu} D_{1}} = \left( {U_{1},V_{1}} \right)} \\ {{{Output}\text{:}\mspace{11mu} D_{2}} = \left( {U_{2},V_{2}} \right)} \\ {\mspace{31mu}{{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},{{\gcd\left( {h,U_{1}} \right)} = 1}}} \\ \left. {1.\mspace{11mu} U_{1}^{\prime}}\leftarrow U_{1}^{2} \right. \\ \left. {2.\mspace{11mu} S}\leftarrow{{{h^{- 1}\left( {f + {hV}_{1} + V_{1}^{2}} \right)}/U_{1}}\mspace{11mu}{mod}\mspace{11mu} U_{1}} \right. \\ \left. {3.\mspace{11mu} V_{1}^{\prime}}\leftarrow{{SU}_{1} + V_{1}} \right. \\ \left. {4.\mspace{11mu} U_{2}^{\prime}}\leftarrow{\left( {f + {hV}_{1}^{\prime} + V_{1}^{\prime\; 2}} \right)/U^{\prime}} \right. \\ \left. {5.\mspace{11mu} U_{2}}\leftarrow{{MakeMonic}\mspace{11mu}\left( U_{2}^{\prime} \right)} \right. \\ \left. {6.\mspace{11mu} V_{2}}\leftarrow{V_{1}^{\prime} + {h\mspace{11mu}{mod}\mspace{11mu} U_{2}}} \right. \\ {{7.\mspace{11mu}{return}\mspace{14mu} D_{2}} = \left( {U_{2},V_{2}} \right)} \end{matrix} & \left\lbrack {{Formula}\mspace{20mu} 24} \right\rbrack \end{matrix}$

In step 6 of Algorithm 1, there is a unique polynomial: k(x)=k ₁ x+k ₀, which satisfies: V ₁ ′+h=(k ₁ x+k ₀)U ₂ +V ₂.

This is transformed as follows: V ₁ ′=h+(k ₁ x+k ₀)U ₂ +V ₂.

Substituting this into the expression that appears in step 4, (f+hV ₁ ′+V ₁′²) results in the following: U ₂ ′U ₁ ′=f+h(kU ₂ +V ₂)+k ² U ₂ ² +V ₂ ²  Expression (1).

In the above expression, since (U₂,V₂) is known, from Expression (1), the relational expression between k and U₁′ can be obtained.

In this case, it is to be noted that U ₂ =k ₁ ² U ₂.

The above equation (1) is expanded and rearranged to yield the following:

$\begin{matrix} {\;{{U_{\; 1}}^{\;\prime} = {x^{\; 4} + {\left( {\left( {{k_{\; 1}h_{\; 2}} + {k_{\; 1}^{\; 2}u_{\; 21}} + 1} \right)/k_{\; 1}^{\; 2}} \right)x^{\; 3}} + {\left( {\left( {{k_{\; 1}h_{\; 1}} + {k_{\; 0}h_{\; 2}} + {k_{\; 1}^{\; 2}u_{\; 20}} + k_{\; 0}^{\; 2} + c_{\; 2}} \right)/k_{\; 1}^{\; 2}} \right)x^{\; 2}} + {\left( {\left( {{k_{\; 1}h_{\; 0}} + {k_{\; 0}h_{\; 1}} + {k_{\; 0}^{\; 2}u_{\; 21}} + c_{\; 1}} \right)/k_{\; 1}^{\; 2}} \right)x} + {\left( {{k_{\; 0}h_{\; 0}} + {k_{\; 0}^{\; 2}u_{\; 20}} + c_{\; 0}} \right)/{k_{\; 1}^{\; 2}.}}}}} & {{Expression}\mspace{14mu}(2)} \end{matrix}$

Here, c ₂ =f ₄ +u ₂₁, c ₁ =f ₃ +h ₂ v ₂₁ +u ₂₁ +c ₂ u ₂₁, and c ₀ =f ₂ +h ₂ v ₂₀ +h ₁ v ₂₁ +v ₂₁ ² +c ₂ u ₂₀ +c ₁ u ₂₁ are satisfied.

Further, from step 1, U ₁ ′=U ₁ ². That is, the following expression holds: U ₁ ′=x ⁴ +u ₁₁ ² x ² +u ₁₀ ²  Expression (3)

A relational expression is derived through comparison between the respective coefficients of Expressions (2) and (3) mentioned above, and halving operation can be computed by solving this relational expression. The algorithm prescribing the above-mentioned procedure is shown below as Algorithm 4 [Algorithm 4 Sketch HEC_HLV].

$\begin{matrix} \begin{matrix} {{Algorithm}\mspace{14mu} 4\mspace{14mu}{Sketch}\mspace{14mu}{HEC\_ HLV}} \\ {{{Input}\text{:}\mspace{11mu} D_{2}} = \left( {U_{2},V_{2}} \right)} \\ {{{Output}\text{:}\mspace{11mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}} \\ {\mspace{31mu}\begin{matrix} {{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},} \\ {{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},} \\ {{{\gcd\left( {h,U_{i}} \right)} = 1},{i = 1},2} \end{matrix}} \\ {{1.\mspace{11mu}{reconstruct}\mspace{14mu} k_{0}},k_{1}} \\ {\mspace{25mu}{\left. {1.1\mspace{14mu} V_{1}^{\prime}}\leftarrow{V_{2} + h + {kU}_{2}} \right.,{{k(x)} = {{k_{1}x} + k_{0}}}}} \\ {\mspace{25mu}\left. {1.2\mspace{14mu} U_{1}^{\prime}}\leftarrow{\left( {f + {hV}_{1}^{\prime} + V_{1}^{\prime\; 2}} \right)/\left( {k_{1}^{2}U_{2}} \right)} \right.} \\ {\mspace{25mu}{{1.3\mspace{14mu}{derive}\mspace{14mu} k_{0}},{{k_{1}\mspace{11mu}{from}\mspace{14mu}{coeff}\;\left( {U_{1}^{\prime},3} \right)} = 0},{{{coeff}\left( {U_{1}^{\prime},1} \right)} = 0}}} \\ {{2.\mspace{11mu}{compute}\mspace{14mu} u_{11}\mspace{11mu}{by}\mspace{14mu}{substituting}\mspace{14mu} k_{0}},{k_{1}\mspace{11mu}{in}\mspace{14mu}{{coeff}\left( {U_{1}^{\prime},2} \right)}}} \\ {{3.\mspace{11mu}{compute}\mspace{14mu} u_{10}\mspace{11mu}{by}\mspace{14mu}{substituting}\mspace{14mu} k_{0}},{k_{1}\mspace{11mu}{in}\mspace{14mu}{{coeff}\left( {U_{1}^{\prime},0} \right)}}} \\ \left. {4.\mspace{11mu} U_{1}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right. \\ \left. {5.\mspace{11mu} V_{1}}\leftarrow{V_{2} + h + {{kU}_{2}\mspace{14mu}{mod}\mspace{11mu} U_{1}}} \right. \\ {{6.\mspace{11mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)} \end{matrix} & \left\lbrack {{Formula}\mspace{20mu} 25} \right\rbrack \end{matrix}$

Specifically, the following relational expressions can be obtained.

[Formula 26] k ₁ h ₂ +k ₁ ² u ₂₁+1=0  Expression (4) k ₁ h ₀ +k ₀ h ₁ +k ₀ ² u ₂₁ +c ₁=0  Expression (5) u ₁₁=√{square root over (k ₁ h ₁ +k ₀ h ₂ +k ₁ ² u ₂₀ +k ₀ ² +c ₂)}/k ₁  Expression (6) u ₁₀=√{square root over (k ₀ h ₀ +k ₀ ² u ₂₀ +c ₀)}/k ₁  Expression (7)

It is necessary to compute the correct k₀, k₁ from these relational expressions. This can be computed using the lemma as described below.

[Lemma 1]

It is assumed that that h(x) is an irreducible polynomial. In this case, there is only one k₁ that satisfies the expressions (4) and (5). Further, the expression (5) has roots only for the correct k₁. Further, there is only one k₀ that allows the computation of the halved divisor D₁ in Algorithm 4. Further, the following expression: xh ₂ +x ² u ₁₁+1=0 has roots only for the correct k₀.

The above-mentioned Lemma 1 was applied to Algorithm 4. The detailed computation method of halving is shown as Algorithm 5a [Algorithm 5a Sketch HEC_HLV] below.

$\begin{matrix} \begin{matrix} {{Algorithm}\mspace{14mu} 5a\;{HEC\_ HLV}} \\ {{{Input}\text{:}\mspace{11mu} D_{2}} = \left( {U_{2},V_{2}} \right)} \\ {{{Output}\text{:}\mspace{11mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}} \\ {\mspace{31mu}\begin{matrix} {{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},} \\ {{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},} \\ {{{\gcd\left( {h,U_{i}} \right)} = 1},{i = 1},2} \end{matrix}} \\ {{{1.\mspace{11mu}{Solve}\mspace{14mu} k_{1}h_{2}} + {k_{1}^{2}u_{21}} + 1} = 0} \\ {\mspace{31mu}{\left. {invu}\leftarrow{1/u_{21}} \right.,\left. {invh}\leftarrow{1/h_{2}^{2}} \right.,\left. \alpha\leftarrow{h_{2}{invu}} \right.,\left. \gamma\leftarrow{u_{21}{invh}} \right.}} \\ {\mspace{31mu}{\left. k_{1}\leftarrow{{H(\gamma)}\alpha} \right.,\left. k_{1}^{\prime}\leftarrow{k_{1} + \alpha} \right.}} \\ {{2.\mspace{11mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{1}}\mspace{14mu}} \\ {\mspace{25mu}\begin{matrix} {{{{by}\mspace{14mu}{solving}\mspace{14mu} k_{1}h_{0}} + {k_{0}h_{1}} + {k_{0}^{2}u_{21}} + c_{1}} = 0} \\ {\left. c_{2}\leftarrow{f_{4} + u_{21}} \right.,\left. c_{1}\leftarrow{f_{3} + {h_{2}v_{21}} + u_{20} + {c_{2}u_{21}}} \right.} \end{matrix}} \\ {\mspace{31mu}\left. c_{0}\leftarrow{f_{2} + {h_{2}v_{20}} + {h_{1}v_{21}} + v_{21}^{2} + {c_{2}u_{20}} + {c_{1}u_{21}}} \right.} \\ {\mspace{31mu}{\left. \alpha\leftarrow{h_{1}{invu}} \right.,\left. w\leftarrow{u_{21}/h_{1}^{2}} \right.,\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right)w} \right.}} \\ {\mspace{31mu}{{{{if}\mspace{14mu}{{Tr}(\gamma)}} = \left. {1\mspace{14mu}{then}\mspace{14mu} k_{1}}\leftarrow k_{1}^{\prime} \right.},\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right)w} \right.}} \\ {\mspace{31mu}{\left. k_{0}\leftarrow{{H(\gamma)}\alpha} \right.,\left. k_{0}^{\prime}\leftarrow{k_{0} + \alpha} \right.}} \\ {{3.\mspace{11mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{0}}\mspace{14mu}} \\ {\mspace{25mu}\begin{matrix} {{{{by}\mspace{14mu}{checking}\mspace{14mu}{trace}\mspace{14mu}{of}\mspace{14mu}{xh}_{2}} + {x^{2}u_{11}} + 1} = 0} \\ {\;{\left. {invk}\leftarrow{1/k_{1}} \right.,}} \\ {\left. u_{11}\leftarrow{{invk}\sqrt{{k_{1}\left( {h_{1} + {k_{1}u_{20}}} \right)} + {k_{0}\left( {h_{2} + k_{0}} \right)} + c_{2}}} \right.,} \\ \left. \gamma\leftarrow{u_{11}{invh}} \right. \end{matrix}} \\ {\mspace{31mu}{{{if}\mspace{14mu}{{Tr}(\gamma)}} = {1\mspace{14mu}{then}}}} \\ {\mspace{95mu}{\left. k_{0}\leftarrow k_{0}^{\prime} \right.,\left. u_{11}\leftarrow{u_{11} + {{invk}\sqrt{\alpha\left( {h_{2} + \alpha} \right)}}} \right.}} \\ {4.\mspace{11mu}{Compute}\mspace{14mu} U_{1}} \\ {\mspace{31mu}\left. u_{10}\leftarrow{{invk}\sqrt{{k_{0}\left( {h_{0} + {k_{0}u_{20}}} \right)} + c_{0}}} \right.} \\ {{5.\mspace{11mu}{Compute}\mspace{14mu} V_{1}} = {V_{2} + h + {{kU}_{2}\mspace{11mu}{mod}\mspace{11mu} U_{1}}}} \\ {\mspace{25mu}\left. w\leftarrow{h_{2} + {k_{1}\left( {u_{11} + u_{21}} \right)} + k_{0}} \right.} \\ {\mspace{25mu}\left. v_{11}\leftarrow{v_{21} + h_{1} + {k_{1}\left( {u_{10} + u_{20}} \right)} + {k_{0}u_{21}} + {u_{11}w}} \right.} \\ {\mspace{25mu}\left. v_{10}\leftarrow{v_{20} + h_{0} + {k_{0}u_{20}} + w} \right.} \\ {\left. {6.\mspace{11mu}{U_{1}(x)}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,\left. {V_{1}(x)}\leftarrow{{v_{11}x} + v_{10}} \right.} \\ {{7.\mspace{11mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)} \end{matrix} & \left\lbrack {{Formula}\mspace{20mu} 27} \right\rbrack \end{matrix}$

In Algorithm 5a [Algorithm 5a] mentioned above, if k₁′, k₀′ are correct values (in other words, if k₁, k₀ are not correct values), the complexity of the algorithm is: 29M+1S+4I+3SR+2H+2T.

Here, M, S, SR, H, and T mean multiplication, squaring, inversion, square root operation, half-trace (operation to find the root of a quadratic equation), and trace (determination as to whether roots exist for a quadratic equation), respectively. The complexity becomes the largest if these k₁′, k₀′ are correct values.

Next, if k₁, k₀ are correct values (in other words, if k₁′, k₀′ are not correct values), the complexity becomes the smallest, so the complexity can be reduced by 2M in step 2, and the complexity can be reduced by 2M+1SR in step 3. That is, the complexity in this case is 25M+1S+4I+2SR+2H+2T, and the complexity becomes the smallest.

Next, if k₁, k₀′ are correct values (in other words, if k₁′, k₀ are not correct values), the complexity can be reduced by 2M+1SR in step 3. That is, the complexity in this case becomes: 27M+1S+4I+2SR+2H+2T.

Lastly, if k₁′, k₀ are correct values (in other words, if k₁, k₀′ are not correct values), the complexity can be reduced by 2M in step 2. That is, the complexity in this case becomes: 27M+1S+4I+3SR+2H+2T.

Upon checking the probabilities with which the above-mentioned four cases occur by computer experiment, it was confirmed that they occur at substantially the same ratio. In the description that follows, it is assumed that the probabilities with which the above-mentioned four cases occur are substantially equal. The averaging of the complexities in the above-mentioned four cases yields 27M+1S+4I+2.5SR+2H+2T.

Next, the exceptional cases:

ExHEC_HLV^(2→1+1);

ExHEC_HLV^(1→2+2); and

ExHEC_HLV^(2→2+2)

are considered. Since the probabilities with which these exceptional cases occur are so low as to be negligible, no evaluation on complexity will be made.

It should be noted that the computation algorithms for these exceptional cases are of the same processing as those described with reference to the flowcharts shown in FIGS. 4 to 6 in the section of (Processing Example 1: Proposed Method A1) described above, that is:

For the algorithm of ExHEC_HLV^(2→1+1), the flowchart shown in FIG. 4;

For the algorithm of ExHEC_HLV^(1→2+2), the flowchart shown in FIG. 5; and

For the algorithm of ExHEC_HLV^(2→2+2), the flowchart shown in FIG. 6.

Further, as for the computation procedure for ExHEC_HLV^(2→2+2) as well, the processing is the same as that of Algorithm 6 [Algorithm 6] described above in the section of (Processing Example 1: Proposed Method A1), and as for the computation procedure for ExHEC_HLV^(1→2+2) as well, the processing is the same as that of Algorithm 7 [Algorithm 7] described above in the section of (Processing Example 1: Proposed Method A1).

Processing Example 8 Proposed Method F2

Processing Example 8 (Proposed Method F2A) relates to a technique aimed at a further increase in the operation speed of the processing example mentioned above (Processing Example 2 Proposed Method F1), which includes computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+h₁x+h₀, f₄=0

A close look at Algorithm 5a applied to above-mentioned processing example (processing Example 7: Proposed Method A2) will reveal that Algorithm 5a contains a large number of multiplication operations by a coefficient h(x) and inversion operations of the coefficient h(x). This means that the complexities of multiplication and inversion operations can be reduced by manipulating the coefficient h(x). It should be noted that according to the document (Non-patent Document 19: T. Lange. Efficient Doubling on Genus Two Curves over Binary Fields, SAC 2004, pre-proceedings, pp. 189-202, 2004.), h₂=1, f₄=0 are used to achieve fast computation. The complexity of HarleyDBL in the case where these parameters are used is 21M+5S+1I.

While the conditions for Processing Example 8 (Proposed Method F2) described here are also set in conformity with those mentioned above, since an irreducible polynomial is assumed for h(x) due to Lemma 1, h(x)=x ² +h ₁ x+h ₀, and Tr(h ₀ /h ₁ ²)=1 are set (the necessary and sufficient condition for the quadratic equation ax²+bx+c=0 to be an irreducible polynomial is Tr(ac/b²)=1). The computation method in this case is shown in Algorithm 8 [Algorithm 8] HEC_HLV(h₂=1, f₄=0).

$\begin{matrix} \begin{matrix} {{Algorithm}\mspace{14mu} 8a\mspace{11mu}{HEC\_ HLV}\left( {{h_{2} = 1},{f_{4} = 0}} \right)} \\ {{{{Input}\text{:}\mspace{11mu} D_{2}} = \left( {U_{2},V_{2}} \right)},{1/h_{1}^{2}}} \\ {{{Output}\text{:}\mspace{11mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}} \\ {\mspace{31mu}\begin{matrix} {{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},} \\ {{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},} \\ {{{\gcd\left( {h,U_{i}} \right)} = 1},{i = 1},2} \end{matrix}} \\ {{{1.\mspace{11mu}{Solve}\mspace{14mu} k_{1}} + {k_{1}^{2}u_{21}} + 1} = 0} \\ {\mspace{25mu}{\left. \alpha\leftarrow{1/u_{21}} \right.,\left. k_{1}\leftarrow{{H\left( u_{21} \right)}\alpha} \right.,\left. k_{1}^{\prime}\leftarrow{k_{1} + \alpha} \right.}} \\ {{{2.\mspace{11mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{1}\mspace{11mu}{by}\mspace{14mu}{solving}\mspace{14mu} k_{1}h_{0}} + {k_{0}h_{1}} + {k_{0}^{2}u_{21}} + c_{1}} = 0} \\ {\mspace{25mu}\left. c_{1}\leftarrow{f_{3} + v_{21} + u_{20} + u_{21}^{2}} \right.} \\ {\mspace{25mu}\left. c_{0}\leftarrow{f_{2} + v_{20} + {v_{21}\left( {h_{1} + v_{21}} \right)} + {u_{21}\left( {u_{20} + c_{1}} \right)}} \right.} \\ {\mspace{25mu}{\left. w_{0}\leftarrow{u_{21}/h_{1}^{2}} \right.,\left. \alpha\leftarrow{h_{1}\alpha} \right.,\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right)w_{0}} \right.}} \\ {\mspace{25mu}{{{{if}\mspace{14mu}{{Tr}(\gamma)}} = \left. {1\mspace{14mu}{then}\mspace{14mu} k_{1}}\leftarrow k_{1}^{\prime} \right.},\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right)w_{0}} \right.}} \\ {\mspace{25mu}{\left. k_{0}\leftarrow{{H(\gamma)}\alpha} \right.,\left. k_{0}^{\prime}\leftarrow{k_{0} + \alpha} \right.}} \\ {{{3.\mspace{11mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{0}\mspace{14mu}{by}\mspace{14mu}{checking}\mspace{14mu}{trace}\mspace{14mu}{of}\mspace{14mu} x} + {x^{2}u_{11}} + 1} = 0} \\ {\mspace{25mu}{\left. w_{0}\leftarrow k_{1}^{2} \right.,\left. w_{1}\leftarrow{{w_{0}u_{20}} + {k_{1}h_{1}} + u_{21}} \right.}} \\ {\mspace{25mu}{\left. w_{2}\leftarrow{k_{0} + \sqrt{w_{1} + k_{0}}} \right.,\left. w_{4}\leftarrow{{k_{1}u_{21}} + 1} \right.,\left. u_{11}\leftarrow{w_{2}w_{4}} \right.}} \\ {\mspace{25mu}{{{if}\mspace{14mu}{{Tr}\left( u_{11} \right)}} = {1\mspace{14mu}{then}}}} \\ {\mspace{104mu}{\left. k_{0}\leftarrow k_{0}^{\prime} \right.,\left. w_{2}\leftarrow{k_{0} + \sqrt{w_{1} + k_{0}}} \right.,\left. u_{11}\leftarrow{w_{2}w_{4}} \right.}} \\ {4.\mspace{11mu}{Compute}\mspace{14mu} U_{1}} \\ {\mspace{25mu}{\left. w_{1}\leftarrow{k_{0}u_{20}} \right.,\left. w_{5}\leftarrow{w_{4} + 1} \right.,\left. w_{6}\leftarrow{\left( {k_{0} + k_{1}} \right)\left( {u_{20} + u_{21}} \right)} \right.}} \\ {\mspace{25mu}\left. u_{10}\leftarrow{w_{4}\sqrt{{k_{0}\left( {w_{1} + h_{0}} \right)} + c_{0}}} \right.} \\ {{5.\mspace{11mu}{Compute}\mspace{14mu} V_{1}} = {V_{2} + h + {{kU}_{2}\mspace{11mu}{mod}\mspace{11mu} U_{1}}}} \\ {\mspace{25mu}{\left. w_{4}\leftarrow{w_{5} + k_{0} + 1} \right.,\left. w_{5}\leftarrow{w_{1} + w_{5} + w_{6} + v_{21} + h_{1}} \right.}} \\ {\mspace{25mu}{\left. w_{6}\leftarrow{w_{1} + v_{20} + h_{0}} \right.,\left. w_{7}\leftarrow{w_{2} + w_{4}} \right.}} \\ {\mspace{25mu}{\left. w_{1}\leftarrow{w_{7}u_{10}} \right.,\left. w_{3}\leftarrow{\left( {k_{1} + w_{7}} \right)\left( {u_{10} + u_{11}} \right)} \right.}} \\ {\mspace{25mu}\left. v_{11}\leftarrow{w_{1} + w_{2} + w_{3} + w_{5}} \right.} \\ {\mspace{25mu}\left. v_{10}\leftarrow{w_{1} + w_{6}} \right.} \\ {\left. {6.\mspace{11mu}{U_{1}(x)}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,\left. {V_{1}(x)}\leftarrow{{v_{11}x} + v_{10}} \right.} \\ {{7.\mspace{11mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)} \end{matrix} & \left\lbrack {{Formula}\mspace{20mu} 28} \right\rbrack \end{matrix}$

Here, a method of reducing the number of inversion operations is considered. While Algorithm 5a [Algorithm 5a] requires four inversion operations of 1/u₂₁, 1/h₂ ², 1/h₁ ², 1/k₁, here, by setting as h₂=1, the number of required inversion operations can be reduced to three. Further, since h₁ is a curve parameter, by computing 1/h₁ ² and giving this as an input in advance, the number of required inversion operations can be reduced to two of 1/u₂₁ and 1/k₁.

Further, with regard to 1/k₁, this can be found by one multiplication operation and one addition operation on the basis of Expression 4 described with reference to [Processing Example 7: Proposed Method A2] mentioned above, that is, 1/k ₁ =h ₂ +k ₁ u ₂₁. Due to these operations, the number of inversion operations required for Algorithm 8a [Algorithm 8a] HEC_HLV(h₂=1, f₄=0) mentioned above is only one, 1/u₂₁.

As a result, the complexity of Algorithm 8a [Algorithm 8a] HEC_HLV(h₂=1, f₄=0) mentioned above is as follows.

(a) If k₁, k₀ are correct values: 18M+2S+1I+2SR+2H+2T

(b) If k₁, k₀′ are correct values: 19M+2S+1I+3SR+2H+2T

(c) If k₁′, k₀ are correct values: 20M+2S+1I+2SR+2H+2T

(d) If k₁′, k₀′ are correct values: 21M+2S+1I+3SR+2H+2T

The averaging of all of the above-mentioned cases (a) to (d) yields 19.5 M+2S+1I+2.5SR+2H+2T.

The complexity of HarleyDBL was 21M+5S+1I. Here, according to the document [(Non-patent Document 15) E. Knudsen. Elliptic Scalar Multiplication Using Point Halving. ASIACRYPTO '99, LNCS 1716, pp. 135-149, Springer-Verlag, 1999.], it is known that when a finite field is defined by a normal basis, the complexities of S (squaring), SR (square root operation), H (half-trace) (operation to find the root of a quadratic equation)), and T (trace (determination as to whether roots exist for a quadratic equation)) can be ignored, and only the complexities of M (multiplication) and I (inversion) need to be taken into account. Therefore, when using a normal basis, Algorithm 8a is faster than HarleyDBL by 1.5M.

Further, when a finite field is defined by a polynomial basis, according to the document [(Non-patent Document 16) K. Fong, D. Hankerson, J. Lopez, and A. Menezes. Field inversion operation and point halving revised. Technical Report CORR2003-18, www.cacr.math.uwaterloo.ca/techreports/2003/corr2003-18.pdf18], it is known that in comparison to the complexity of M (multiplication), generally, the complexities of SR (square root operation) and H (half-trace) (operation to find the root of a quadratic equation)) are about SR=H=0.5M. Further, the complexity of T (trace (determination as to whether roots exist for a quadratic equation)) can be ignored. Further, the complexity of S (squaring) is known to be only about several tenths of M (multiplication). However, it is also known that depending on the way in which the polynomial basis is chosen, the complexity of SR may become less than 0.5M. It should be noted that exceptional cases can be computed on the basis of the exceptional cases in Processing Example 7 (proposed Method A2) described above.

Processing Example 9 Proposed Method B2

Processing Example 9 (Proposed Method B2) relates to a technique aimed at a further increase in the operation speed of Processing Example mentioned above (Processing Example 3: Proposed Method B1), which includes computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+x+h₀, f₄=0

As has been described with reference to Processing Example 8 (Proposed Method F2) mentioned above, a close look at the computation algorithm for halving described with reference to Processing Example 7 (Proposed Method A2), that is, Algorithm 5a [Algorithm 5a HEC_HLV] will reveal that Algorithm 5a contains a large number of multiplication operations by a coefficient h(x) and inversion operations of the coefficient h(x). This means that the complexities of multiplications and inversion operations can be reduced by manipulating the coefficient h(x). In the document [J. Pelzl, T. Wollinger, J. Guajardo, and C. Paar. Hyperelliptic curve Cryptosystems: Closing the Performance Gap to Elliptic Curves. Cryptology ePrint Archive, 2003/026, IACR, 2003], there is disclosed an example in which h₂, h₁ε{0, 1}, f₄=0 is used to achieve fast computation.

The complexity of HarleyDBL in the case where these parameters are used is 18M+7S+1I.

While the conditions for Processing Example 9 (Proposed Method B2) are also set in conformity with those mentioned above, since an irreducible polynomial is assumed for h(x) due to Lemma 1 mentioned above, h(x)=x ² +x+h ₀, and Tr(h ₀)=1 are set (the necessary and sufficient condition for the quadratic equation ax²+bx+c=0 to be an irreducible polynomial is Tr(ac/b²)=1)

The computation method in this case is shown below as Algorithm 10a [Algorithm 10a HEC_HLV(h₂=h₁=1, f₄=0).

$\begin{matrix} \begin{matrix} {{Algorithm}\mspace{14mu} 10a\mspace{11mu}{HEC\_ HLV}\left( {{h_{2} = {h_{1} = 1}},{f_{4} = 0}} \right)} \\ {{{Input}\text{:}\mspace{11mu} D_{2}} = \left( {U_{2},V_{2}} \right)} \\ {{{Output}\text{:}\mspace{11mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}} \\ {\mspace{31mu}\begin{matrix} {{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},} \\ {{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},} \\ {{{\gcd\left( {h,U_{i}} \right)} = 1},{i = 1},2} \end{matrix}} \\ {{{1.\mspace{11mu}{Solve}\mspace{14mu} k_{1}} + {k_{1}^{2}u_{21}} + 1} = 0} \\ {\mspace{25mu}{\left. \alpha\leftarrow{1/u_{21}} \right.,\left. k_{1}\leftarrow{{H\left( u_{21} \right)}\alpha} \right.,\left. k_{1}^{\prime}\leftarrow{k_{1} + \alpha} \right.}} \\ {{{2.\mspace{11mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{1}\mspace{14mu}{by}\mspace{14mu}{solving}\mspace{14mu} k_{1}h_{0}} + k_{0} + {k_{0}^{2}u_{21}} + c_{1}} = 0} \\ {\mspace{25mu}\left. c_{1}\leftarrow{f_{3} + v_{21} + u_{20} + u_{21}^{2}} \right.} \\ {\mspace{25mu}\left. c_{0}\leftarrow{f_{2} + v_{20} + v_{21} + v_{21}^{2} + {u_{21}\left( {u_{20} + c_{1}} \right)}} \right.} \\ {\mspace{25mu}\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right)u_{21}} \right.} \\ {\mspace{25mu}{{{{if}\mspace{14mu}{{Tr}(\gamma)}} = \left. {1\mspace{14mu}{then}\mspace{14mu} k_{1}}\leftarrow k_{1}^{\prime} \right.},\left. \gamma\leftarrow{\gamma + h_{0}} \right.}} \\ {\mspace{25mu}{\left. k_{0}\leftarrow{{H(\gamma)}\alpha} \right.,\left. k_{0}^{\prime}\leftarrow{k_{0} + \alpha} \right.}} \\ {{{3.\mspace{11mu}{Select}\mspace{14mu}{correct}\mspace{14mu} k_{0}\mspace{14mu}{by}\mspace{14mu}{checking}\mspace{14mu}{trace}\mspace{14mu}{of}\mspace{14mu} x} + {x^{2}u_{11}} + 1} = 0} \\ {\mspace{25mu}{\left. w_{0}\leftarrow k_{1}^{2} \right.,\left. w_{1}\leftarrow{{w_{0}u_{20}} + k_{1} + u_{21}} \right.}} \\ {\mspace{25mu}{\left. w_{2}\leftarrow{k_{0} + \sqrt{w_{1} + k_{0}}} \right.,\left. w_{4}\leftarrow{{k_{1}u_{21}} + 1} \right.,\left. u_{11}\leftarrow{w_{2}w_{4}} \right.}} \\ {\mspace{25mu}{{{if}\mspace{14mu}{{Tr}\left( u_{11} \right)}} = {1\mspace{14mu}{then}}}} \\ {\mspace{104mu}{\left. k_{0}\leftarrow k_{0}^{\prime} \right.,\left. w_{2}\leftarrow{k_{0} + \sqrt{w_{1} + k_{0}}} \right.,\left. u_{11}\leftarrow{w_{2}w_{4}} \right.}} \\ {4.\mspace{11mu}{Compute}\mspace{14mu} U_{1}} \\ {\mspace{25mu}{\left. w_{1}\leftarrow{k_{0}u_{20}} \right.,\left. w_{5}\leftarrow{w_{4} + 1} \right.,\left. w_{6}\leftarrow{\left( {k_{0} + k_{1}} \right)\left( {u_{20} + u_{21}} \right)} \right.}} \\ {\mspace{25mu}\left. u_{10}\leftarrow{w_{4}\sqrt{{k_{0}\left( {w_{1} + h_{0}} \right)} + c_{0}}} \right.} \\ {{5.\mspace{11mu}{Compute}\mspace{14mu} V_{1}} = {V_{2} + h + {{kU}_{2}\mspace{11mu}{mod}\mspace{11mu} U_{1}}}} \\ {\mspace{25mu}{\left. w_{4}\leftarrow{w_{5} + k_{0} + 1} \right.,\left. w_{5}\leftarrow{w_{1} + w_{5} + w_{6} + v_{21} + 1} \right.}} \\ {\mspace{25mu}{\left. w_{6}\leftarrow{w_{1} + v_{20} + h_{0}} \right.,\left. w_{7}\leftarrow{w_{2} + w_{4}} \right.}} \\ {\mspace{25mu}{\left. w_{1}\leftarrow{w_{7}u_{10}} \right.,\left. w_{3}\leftarrow{\left( {k_{0} + w_{7}} \right)\left( {u_{10} + u_{11}} \right)} \right.}} \\ {\mspace{25mu}\left. v_{11}\leftarrow{w_{1} + w_{2} + w_{3} + w_{5}} \right.} \\ {\mspace{25mu}\left. v_{10}\leftarrow{w_{1} + w_{6}} \right.} \\ {\left. {6.\mspace{11mu}{U_{1}(x)}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,\left. {V_{1}(x)}\leftarrow{{v_{11}x} + v_{10}} \right.} \\ {{7.\mspace{11mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)} \end{matrix} & \left\lbrack {{Formula}\mspace{20mu} 29} \right\rbrack \end{matrix}$

The complexity of Algorithm 10a [Algorithm 10a HEC_HLV(h₂=h₁=1, f₄=0) mentioned above is as follows.

(a) If k₁, k₀ are correct values: 14M+3S+1I+2SR+2H+2T

(b) If k₁, k₀′ are correct values: 15M+3S+1I+3SR+2H+2T

(c) If k₁′, k₀ are correct values: 14M+3S+1I+2SR+2H+2T

(d) If k₁′, k₀′ are correct values: 15M+3S+1I+3SR+2H+2T

The averaging of all of the above-mentioned cases (a) to (d) yields: 14.5M+3S+1I+3SR+2H+2T. The complexity of HarleyDBL was 18M+7S+1I. Here, as described above, according to the document [E. Knudsen. Elliptic Scalar Multiplication Using Point Halving. ASIACRYPTO '99, LNCS 1716, pp. 135-149, Springer-Verlag, 1999.], it is known that when a finite field is defined by a normal basis, the complexities of S (squaring), SR (square root operation), H (half-trace) (operation to find the root of a quadratic equation)), and T (trace (determination as to whether roots exist for a quadratic equation)) can be ignored, and only the complexities of M (multiplication) and I (inversion) need to be taken into account.

Therefore, when using a normal basis, Algorithm 10a [Algorithm 10a] described above is faster than the conventional algorithm [HarleyDBL] by 3.5M. Further, when a finite field is defined by a polynomial basis, according to the document [K. Fong, D. Hankerson, J. Lopez, and A. Menezes. Field inversion operation and point halving revised. Technical Report CORR2003-18, www.cacr.math.uwaterloo.ca/techreports/2003/corr2003-18.pdf18], it is known that generally, the complexities of SR and H are about SR=H=0.5M. Further, the complexity of T can be ignored. Further, it is known that the complexity of S is only about several tenths of M. However, it is also known that depending on the way in which the polynomial basis is chosen, the complexity of SR may become less than 0.5M.

The curve of Algorithm 10a [Algorithm 10a] mentioned above is also subject to the constraint h₀=1. Since Algorithm 10a [Algorithm 10a] mentioned above involves one multiplication operation of h₀, by setting as h₀=1, the complexity can be reduced by 1M. The complexity found by the averaging of all of the above-mentioned cases (a) to (d) is 13.5M+3S+1I+2.5SR+2H+2T. On the other hand, the complexity of HarleyDBL is 15M+7S+11. It should be noted that when a finite field is defined by a normal basis, the complexities of S, SR, H, T can be ignored, and when a normal basis is used, Algorithm 10a [Algorithm 10a] becomes faster than the conventional algorithm [HarleyDBL] by 1.5M. It should be noted that exceptional cases can be computed on the basis of the exceptional cases in Processing Example 7 (proposed Method A2) described above.

Processing Example 10 Proposed Method C2

Processing Example 10 (Proposed Method C2) relates to a technique aimed at a further increase in the operation speed of the processing example mentioned above (Processing Example 5: Proposed Method C1). That is, when computing the halving of a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with random parameters, a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+h₁x+h₀, f₄=0, and a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+x+h₀, f₄=0, two candidates of the halved value arise. In this case, it is necessary to select the one with the correct value from the two candidates. When selecting the correct one, it is necessary to compute the trace, multiplication, and square root of a finite field. Which one of the two candidates is correct depends on the divisor D. Hence, if the divisor D is fixed, information as to which one of the two candidates is correct is retained in a table in advance, and this table is looked up when selecting the correct value, thereby omitting the above-mentioned extra computations.

Which one of k₁, k₁′ (k₀, k₀′) is correct depends on the input divisor D. Accordingly, if D is fixed, for example, when the base point is previously determined as in the case of Phase 1 of ECDH key exchange, ECDSA signature generation or verification, or the like, [½^(i)]D is computed and information as to which of k₁, k₁′ (k₀, k₀′) is correct is recorded in a table in advance.

For example, two tables T₁, T₀ of the same bit size as the order of the base point are prepared, and the binary expression of these tables is defined as: T ₁=(t _(1,r-1) , . . . , t _(1,0)), and T ₀=(t _(0,r-1) , . . . , t _(0,0))

When finding [½^(i)]D, if such information that if k₁ is correct, then t_(1,i)=0 or else if k₁ is correct, then t_(1,i)=1; and if k₁ is correct, then t_(1,i)=0 or else if k₀′ is correct, then t_(0,i)=1 is stored in the tables, a bit string only about twice the size of the order of the base point suffices as the table size. By looking up these tables, the complexity of halving can be reduced.

The above-mentioned method as applied to Algorithm 8a [Algorithm 8a] HEC_HLV(h₂=1, f₄=0) is represented as Algorithm 9a [Algorithm 9a] HEC_HLV(h₂=1, f₄=0, with table-lookup). The complexity of the algorithm is 18M+2SR+1I+2SR+2H.

$\begin{matrix} {{{Algorithm}\mspace{14mu} 9a\mspace{14mu}{HEC\_ HLV}\left( {{h_{2} = 1},{f_{4} = 0},\;{{table}\text{-}{lookup}}} \right)}{{{{Input}\text{:}\mspace{11mu} D_{2}} = \left( {U_{2},V_{2}} \right)},{1/h_{1}^{2}},t_{0},t_{1}}{{{Output}\text{:}\mspace{11mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {\left\lbrack {1/2} \right\rbrack D_{2}}}}{{{U_{i}(x)} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}},{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},{{\gcd\left( {h,U_{i}} \right)} = 1},{i = 1},2}{{{1.\mspace{14mu}{Solve}\mspace{14mu} k_{1}} + {k_{1}^{2}u_{21}} + 1} = 0}\left. \alpha\leftarrow{1/u_{21}} \right.{{{if}\mspace{14mu} t_{1}} = \left. {0\mspace{14mu}{then}\mspace{14mu} k_{1}}\leftarrow{{H\left( u_{21} \right)}\alpha\mspace{14mu}{else}\mspace{14mu} k_{1}}\leftarrow{\left( {{H\left( u_{21} \right)} + 1} \right)\alpha} \right.}{{{2.\mspace{14mu}{Solve}\mspace{14mu} k_{1}h_{0}} + {k_{0}h_{1}} + {k_{0}^{2}u_{21}} + c_{1}} = 0}\left. c_{1}\leftarrow{f_{3} + v_{21} + u_{20} + u_{21}^{2}} \right.\left. c_{0}\leftarrow{f_{2} + v_{20} + {v_{21}\left( {h_{1} + v_{21}} \right)} + {u_{21}\left( {u_{20} + c_{1}} \right)}} \right.{\left. w_{0}\leftarrow{u_{21}/h_{1}^{2}} \right.,\left. \alpha\leftarrow{h_{1}\alpha} \right.,\left. \gamma\leftarrow{\left( {c_{1} + {k_{1}h_{0}}} \right)w_{0}} \right.}{{{if}\mspace{14mu} t_{0}} = \left. {0\mspace{14mu}{then}\mspace{14mu} k_{0}}\leftarrow{{H(\gamma)}\alpha\mspace{14mu}{else}\mspace{14mu} k_{0}}\leftarrow{\left( {{H(\gamma)} + 1} \right)\alpha} \right.}{3.\mspace{14mu}{Compute}\mspace{14mu} U_{1}}{\left. w_{0}\leftarrow k_{1}^{2} \right.,\left. w_{1}\leftarrow{{w_{0}u_{20}} + {k_{1}h_{1}} + u_{21}} \right.,\left. w_{2}\leftarrow{k_{0} + \sqrt{w_{1} + k_{0}}} \right.}{\left. w_{4}\leftarrow{{k_{1}u_{21}} + 1} \right.,\left. u_{11}\leftarrow{w_{2}w_{4}} \right.}{\left. w_{1}\leftarrow{k_{0}u_{20}} \right.,\left. w_{5}\leftarrow{w_{4} + 1} \right.,\left. w_{6}\leftarrow{\left( {k_{0} + k_{1}} \right)\left( {u_{20} + u_{21}} \right)} \right.}\left. u_{10}\leftarrow{w_{4}\sqrt{{k_{0}\left( {w_{1} + h_{0}} \right)} + c_{0}}} \right.{{4.\mspace{14mu}{Compute}\mspace{14mu} V_{1}} = {V_{2} + h + {{kU}_{2}\;{mod}\; U_{1}}}}{\left. w_{4}\leftarrow{w_{5} + k_{0} + 1} \right.,\left. w_{5}\leftarrow{w_{1} + w_{5} + w_{6} + v_{21} + h_{1}} \right.}{\left. w_{6}\leftarrow{w_{1} + v_{20} + h_{0}} \right.,\left. w_{7}\leftarrow{w_{2} + w_{4}} \right.}{\left. w_{1}\leftarrow{w_{7}u_{10}} \right.,\left. w_{3}\leftarrow{\left( {k_{1} + w_{7}} \right)\left( {u_{10} + u_{11}} \right)} \right.}\left. v_{11}\leftarrow{w_{1} + w_{2} + w_{3} + w_{5}} \right.\left. v_{10}\leftarrow{w_{1} + w_{6}} \right.{\left. {6.\mspace{14mu}{U_{1}(x)}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,\left. {V_{1}(x)}\leftarrow{{v_{11}x} + v_{10}} \right.}{{7.\mspace{14mu}{return}\mspace{14mu} D_{1}} = \left( {U_{1},V_{1}} \right)}} & \left\lbrack {{Formula}\mspace{14mu} 30} \right\rbrack \end{matrix}$

Specifically, the above-mentioned method as applied to Algorithm 10a [Algorithm 10a HEC_HLV(h₂=h₁=1, f₄=0, with table-lookup)] described above is represented as Algorithm 11a [Algorithm 11a HEC_HLV(h₂=h₁=1, f₄=0, with table-lookup) below.

$\begin{matrix} {{{{Algorithm}\mspace{14mu} 11\; a\mspace{14mu}{HEC\_ HLV}\left( {{h_{2} = {h_{1} = 1}},{f_{4} = 0},{{{with}\mspace{14mu}{table}} - {lookup}}} \right)}{{{Input}\text{:}\mspace{14mu} D_{2}} = \left( {U_{2},V_{2}} \right)},t_{0},{t_{1}{{{Output}\text{:}\mspace{14mu} D_{1}} = {\left( {U_{1},V_{1}} \right) = {{\left\lbrack {1/2} \right\rbrack D_{2}{U_{i}(x)}} = {x^{2} + {u_{i\; 1}x} + u_{i\; 0}}}}}},{{V_{i}(x)} = {{v_{i\; 1}x} + v_{i\; 0}}},{{\gcd\left( {h,U_{\; 1}} \right)} = 1}}{{{1.\mspace{14mu}{Solve}\mspace{14mu} k_{1}} + {k_{1}^{2}u_{21}} + 1} = {0\mspace{14mu}{and}\mspace{14mu}{select}\mspace{14mu}{correct}\mspace{14mu} k_{\; 1}\mspace{14mu}{via}\mspace{14mu} t_{\; 1}}}\left. \alpha\leftarrow{1/u_{21}} \right.{{{if}\mspace{14mu} t_{1}} = \left. {0\mspace{14mu}{then}\mspace{14mu} k_{1}}\leftarrow{{H\left( u_{21} \right)}\alpha\mspace{14mu}{else}\mspace{14mu} k_{\; 1}}\leftarrow{\left( {{H\left( u_{\; 21} \right)} + 1} \right)\alpha} \right.}{{{2.\mspace{14mu}{Solve}\mspace{14mu} k_{1}h_{0}} + k_{0} + {k_{0}^{2}u_{21}} + c_{1}} = {0\mspace{14mu}{and}\mspace{14mu}{select}\mspace{14mu}{correct}\mspace{14mu} k_{\; 0}\mspace{14mu}{via}\mspace{14mu} t_{\; 0}}}{\left. c_{1}\leftarrow{f_{3} + v_{21} + u_{20} + {u_{21}^{2}c_{0\;}}}\leftarrow{f_{2} + v_{20} + v_{21} + v_{21}^{2} + {{u_{21}\left( {u_{20} + c_{1}} \right)}\gamma}}\leftarrow{\left( {c_{1}k_{1}h_{0}} \right)u_{21}k_{0}}\leftarrow{{H(\gamma)}\alpha{if}\mspace{14mu} t_{0}} \right. = \left. {0\mspace{14mu}{then}\mspace{14mu} k_{0}}\leftarrow{{H(\gamma)}\alpha\mspace{14mu}{else}\mspace{14mu} k_{\; 0}}\leftarrow{\left( {{H(\gamma)} + 1} \right)\alpha} \right.}{3.\mspace{14mu}{Compute}\mspace{14mu} U_{1}}{\left. w_{0}\leftarrow k_{1}^{2} \right.,\left. w_{1}\leftarrow{{w_{0}u_{20}} + k_{1} + {u_{21}w_{2}}}\leftarrow{k_{0} + {\sqrt{w_{1} + k_{0}}w_{4}}}\leftarrow{{k_{1}u_{21}} + 1} \right.,\left. u_{11}\leftarrow{w_{2}w_{4}w_{1}}\leftarrow{k_{0}u_{20}} \right.,\left. w_{5\;}\leftarrow{w_{4} + 1} \right.,\left. w_{\; 6}\leftarrow{\left( {k_{\; 0} + k_{\; 1}} \right)\left( {u_{\; 20} + u_{\; 21}} \right)} \right.}{{\left. u_{10}\leftarrow{w_{4}\sqrt{{k_{0}\left( {w_{1} + h_{0}} \right)} + c_{0}}5.\mspace{14mu}{Compute}\mspace{14mu} V_{1}} \right. = \left. {V_{2} + h + {{kU}_{2}\;{mod}\; U_{1}w_{4}}}\leftarrow{w_{5} + k_{0} + 1} \right.},\left. w_{5}\leftarrow{w_{1} + w_{5} + w_{6} + v_{21} + {1w_{6}}}\leftarrow{w_{1} + v_{20} + h_{0}} \right.,\left. w_{7}\leftarrow{w_{2} + {w_{4}w_{1}}}\leftarrow{w_{7}u_{10}} \right.,\left. w_{3}\leftarrow{\left( {k_{0} + w_{7}} \right)\left( {u_{10} + u_{11}} \right)v_{11}}\leftarrow{w_{1} + w_{2} + w_{3} + {w_{5}v_{10}}}\leftarrow{w_{1} + {w_{6}6.\mspace{14mu}{U_{1}(x)}}}\leftarrow{x^{2} + {u_{11}x} + u_{10}} \right.,{\left. {V_{1}(x)}\leftarrow{{v_{11}x} + {v_{10}7.\mspace{14mu}{return}\mspace{14mu} D_{1}}} \right. = \left( {U_{1},V_{1}} \right)}}} & \left\lbrack {{Formula}\mspace{14mu} 31} \right\rbrack \end{matrix}$

The complexity of Algorithm 11a [Algorithm 11a HEC_HLV(h₂=h₁=1, f₄=0, with table-lookup) is 14M+3S+1I+2SR+2H, and further, by setting as h₀=1, the complexity can be reduced by 1M. The complexity in this case becomes 13M+3S+1I+2SR+2H.

Processing Example 11 Proposed Method D2

Processing Example 11 (Proposed Method D2) relates to a method of computing the scalar multiplication of a divisor by using the method of computing the halving of a divisor as set forth in each of Processing Examples 7 to 10.

A method of computing scalar multiplication using the halving of a rational point on an elliptic curve is disclosed in each of the documents [E. Knudsen. Elliptic Scalar Multiplication Using Point Halving. ASIACRYPTO '99, LNCS 1716, pp. 135-149, Springer-Verlag, 1999.] and [K. Fong, D. Hankerson, J. Lopez, and A. Menezes. Field inversion operation and point halving revised. Technical Report CORR2003-18, www.cacr.math.uwaterloo.ca/techreports/2003/corr2003-18.pdf]. A method of computing scalar multiplication using the halving of a divisor on a hyperelliptic curve is executed on the basis of the scalar multiplication disclosed in those documents. Here, it is assumed that the divisor D subject to scalar multiplication is a prime number of a large order. Further, it is assumed that a scalar value d is an integer 0<d<r. To perform scalar multiplication using halving, first, the scalar value d expressed in binary representation needs to be expressed in half representation.

Here, it is assumed that m=└ log ₂ r┘  [Formula 32] Further, the remainder when d is multiplied by 2^(m) and then divided by r, that is,

$\begin{matrix} {\sum\limits_{i = 0}^{m}\;\overset{\;}{\left. {{\hat{d}}_{i}2^{i}}\leftarrow{2^{m}{d\left( {{mod}\mspace{14mu} r} \right)}} \right.}} & \left\lbrack {{Formula}\mspace{14mu} 33} \right\rbrack \end{matrix}$ is found. Next, this is divided by 2^(m) to yield

$\begin{matrix} \left. {\sum\limits_{i = 0}^{m}\;{d_{i}/2^{i}}}\leftarrow{\sum\limits_{i = 0}^{m}\;{{\hat{d}}_{i}{2^{i - m}.}}} \right. & \left\lbrack {{Formula}\mspace{14mu} 34} \right\rbrack \end{matrix}$ The scalar value expressed by

$\begin{matrix} {\sum\limits_{i = 0}^{m}\;{d_{i}/2^{i}}} & \left\lbrack {{Formula}\mspace{14mu} 35} \right\rbrack \end{matrix}$ in the above-mentioned expression is used for scalar multiplication [halve-and-add binary] method using halving. Here, d _(i) ,{circumflex over (d)} _(i)ε{0,1}  [Formula 36]

The halve-and-add binary method (right-to-left) and halve-and-add binary method (left-to-right) are represented below as Algorithm 12a [Algorithm 12a] and Algorithm 13a [Algorithm 13a], respectively.

$\begin{matrix} {{{Algorithm}\mspace{14mu} 12\; a\mspace{14mu}{Halve}\text{-}{and}\text{-}{add}\mspace{14mu}{binary}\mspace{14mu}{{Method}\left( {{right}\text{-}{to}\text{-}{left}} \right)}}{{{{Input}\text{:}\mspace{14mu} D} \in {{{J\left( F_{2^{n}} \right)}\mspace{14mu}{such}\mspace{14mu}{that}\mspace{14mu} 2\; D} \neq O}},{d \in Z},{r\text{:}\mspace{14mu}{order}\mspace{14mu}{of}\mspace{14mu} D},{m = \left\lfloor {\log_{\; 2}r} \right\rfloor}}{{Output}\text{:}\mspace{14mu}{scalar}\mspace{14mu}{multiplication}\mspace{14mu}{dD}}{\left. {1.\mspace{14mu}{\sum\limits_{i = 0}^{m}{{\hat{d}}_{i}2^{i}}}}\leftarrow{n^{m}d\;{mod}\; r} \right.,\;{{\hat{d}}_{i} \in \left\{ {0,1} \right\}}}{\left. {2.\mspace{14mu}{\sum\limits_{i = 0}^{m}{d_{i}/2^{i}}}}\leftarrow{\sum\limits_{i = 0}^{m}{{\hat{d}}_{i}2^{i - m}}} \right.,{d_{i} \in \left\{ {0,1} \right\}}}{\left. {3.\mspace{14mu} Q}\leftarrow O \right.,\left. R\leftarrow D \right.}{4.\mspace{14mu}{for}\mspace{14mu} i\mspace{14mu}{from}\mspace{14mu} 0\mspace{14mu}{to}\mspace{14mu} m\mspace{14mu}{do}\text{:}}{{{if}\mspace{14mu} d_{i}} = \left. {1\mspace{14mu}{then}\mspace{14mu} Q}\leftarrow{Q + R} \right.}\left. R\leftarrow{{HEC\_ HLV}(R)} \right.{5.\mspace{14mu}{return}\mspace{14mu} Q}} & \left\lbrack {{Formula}\mspace{14mu} 37} \right\rbrack \end{matrix}$

$\begin{matrix} {{{{{{Algorithm}\mspace{14mu} 13\; a\mspace{14mu}{Halve}\text{-}{and}\text{-}{add}\mspace{14mu}{binary}\mspace{14mu}{Method}\;\left( {{left}\text{-}{to}\text{-}{right}} \right)}{{{Input}\text{:}\mspace{14mu} D} \in {{{J\left( F_{2^{n}} \right)}\mspace{14mu}{such}\mspace{14mu}{that}\mspace{14mu} 2\; D} \neq O}}},{d \in Z},{r\text{:}\mspace{14mu}{order}\mspace{14mu}{of}\mspace{14mu} D},{m = \left\lfloor {\log_{\; 2}r} \right\rfloor}}{{Output}\text{:}\mspace{14mu}{scalar}\mspace{14mu}{multiplication}\mspace{14mu}{dD}}{\left. {1.\mspace{14mu}{\sum\limits_{i = 0}^{m}{{\hat{d}}_{i}2^{i}}}}\leftarrow{2^{m}d\;{mod}\; r} \right.,{{\hat{d}}_{i} \in \left\{ {0,1} \right\}}}{\left. {2.\mspace{14mu}{\sum\limits_{i = 0}^{m}{d_{i}/2^{i}}}}\leftarrow{\sum\limits_{i = 0}^{m}{{\hat{d}}_{i}2^{i - m}}} \right.,{d_{i} \in \left\{ {0,1} \right\}}}\left. {3.\mspace{14mu} Q}\leftarrow O \right.{4.\mspace{14mu}{for}\mspace{14mu} i\mspace{14mu}{from}\mspace{14mu} m\mspace{14mu}{downto}\mspace{14mu} 0\mspace{14mu}{do}\text{:}}\left. Q\leftarrow{{HEC\_ HLV}(Q)} \right.{{{if}\mspace{14mu} d_{i}} = \left. {1\mspace{14mu}{then}\mspace{14mu} Q}\leftarrow{Q + D} \right.}{5.\mspace{14mu}{return}\mspace{14mu} Q}}\mspace{11mu}} & \left\lbrack {{Formula}\mspace{14mu} 38} \right\rbrack \end{matrix}$

HEC_HLV that appears in step 4 of each of Algorithm 12a [Algorithm 12a] and Algorithm 13a [Algorithm 13a] mentioned above may be HEC_HLV of Algorithm 5a [Algorithm 5a] described above using a random curve, HEC_HLV with constraints h₂=1, f₄=0 provided to the curve parameters of Algorithm 8a [Algorithm 8a], HEC_HLV with constraints h₂=h₁=1, f₄=0 provided to the curve parameters of Algorithm 10a [Algorithm 10a], or HEC_HLV with constraints h₂=h₁=h₀=1, f₄=0 provided to the curve parameters of Algorithm 10a [Algorithm 10a].

In the case of the halving operation descried with reference to Processing Examples 1 to 6 above, for example, 1/u₂₁ and 1/u₁₁ are required for the input and output, respectively, in Algorithm 10 shown in [Processing Example 3 (Proposed Method B1)] or Algorithm 11 shown in [Processing Example 5 (Proposed Method C1)] described above. Accordingly, ½^(i)-times multiplication of the base point D, that is, (½)D,( 1/22)D,( 1/23)D,(½_(i))D . . . can be given with the output 1/u₁₁ of the previous halving operation taken as the input of the next halving operation, thereby enabling efficient computation. Accordingly, when computing a scalar multiple by the halve-and-add binary method (right-to-left), this can be accomplished by adding the ½^(i) multiple of the base point D as appropriate. Scalar multiplication can thus performed in an efficient manner in the case of the right-to-left method.

On the other hand, in the case of the left-to-right method, there are cases where the output of the previous halving operation cannot be taken as the input of the next halving operation. In step 4 of Algorithm 13a [Algorithm 13a] described above, first, an intermediate result Q obtained halfway through the algorithm is multiplied by ½ (Q←HEC_HLV (Q)), and if one bit of the scalar value is 1, the base point is added to the intermediate result (if di=1 then Q←Q+D). Accordingly, if a bit is 1, for the next bit, the output 1/u₁₁ of the previous halving operation cannot be given as the input of the next halving operation, so it is necessary to compute an input value 1/u₂₁ anew. Inversion of a finite field involves much more complexity than multiplication. Therefore, in the case of the halving method described in Processing Examples 1 to 6 above, when the left-to-right is applied, it is necessary to perform an extra inversion operation for generating the input value of halving, which detracts from the efficiency of computation. However, according to the proposed methods described with reference to the processing examples from Processing Example 7 onward, 1/u₂₁ is not required for the input, so computation can be performed with the same complexity irrespective of whether the left-to-right or right-to-left method is employed.

Further, HEC_HLV in step 4 of the halve-and-add binary method (right-to-left) of Algorithm 12a [Algorithm 12a] described above may be HEC_HLV with the table-lookup method applied to Algorithm 5a [Algorithm 5a] using a random curve, HEC_HLV with constraints h₂=1, f₄=0 provided and the table-lookup method applied to the curve parameters of Algorithm 8a [Algorithm 8a], HEC_HLV with constraints h₂=h₁=1, f₄=0 provided and the table-lookup method applied to the curve parameters of Algorithm 10a [Algorithm 10a], or HEC_HLV with constraints h₂=h₁=h₀=1, f₄=0 provided and the table-lookup method applied to the curve parameters of Algorithm 10a [Algorithm 10a].

Further, the window method can be applied other than the binary method. Let D represent the input divisor and w represent the window width. The divisor for which the intermediate result is substituted is represented as Q(O.

With respect to integer i=(i _(w-1) , i _(w-2) . . . i ₀)2ε{0, 1, . . . , 2_(w)−1}, the following preliminary computation:

$\begin{matrix} \left. D_{i}\leftarrow{\sum\limits_{j = 0}^{w - 1}\;{\frac{i_{j}}{2^{j}}D}} \right. & \left\lbrack {{Formula}\mspace{14mu} 39} \right\rbrack \end{matrix}$ is carried out to compute a table composed of 2^(w) divisors in advance.

Further, the scalar value d is expanded in ½^(w)-ary representation as follows.

$\begin{matrix} \left. {\sum\limits_{i = 0}^{l}{c_{i}/\left( 2^{w} \right)^{i}}}\leftarrow{\sum\limits_{i = 0}^{m}{d_{i}/2^{i}}} \right. & \left\lbrack {{Formula}\mspace{14mu} 40} \right\rbrack \end{matrix}$ First, halving is applied to Q for w times to give Q←(½^(w))Q. Next, a scalar value c₁ is scanned from the most significant bits of d with the window width w, and a corresponding divisor value in the table is looked up, and this is added to the result as follows. Q←Q+D_(c1) This is repeated down to c₀.

This computation method [halve-and-add window method] is represented below as Algorithm 14a [Algorithm 14a].

$\begin{matrix} {{{{Algorithm}\mspace{14mu} 14\; a\mspace{14mu}{Halve}\text{-}{and}\text{-}{add}\mspace{14mu}{Window}\mspace{14mu}{Method}}{{{Input}\text{:}\mspace{14mu} D} \in {{{J\left( F_{2^{n}} \right)}\mspace{14mu}{such}\mspace{14mu}{that}\mspace{14mu} 2\; D} \neq O}},{d \in Z},{r\text{:}\mspace{14mu}{order}\mspace{14mu}{of}\mspace{14mu} D},{m = \left\lfloor {\log_{\; 2}r} \right\rfloor}}{{Output}\text{:}\mspace{14mu}{scalar}\mspace{14mu}{multiplication}\mspace{14mu}{dD}}{\left. {1.\mspace{14mu}{\sum\limits_{i = 0}^{m}{{\hat{d}}_{i}2^{i}}}}\leftarrow{2^{m}d\;{mod}\; r} \right.,{{\hat{d}}_{i} \in \left\{ {0,1} \right\}}}{\left. {2.\mspace{14mu}{\sum\limits_{i = 0}^{m}{d_{i}/2^{i}}}}\leftarrow{\sum\limits_{i = 0}^{m}{{\hat{d}}_{i}2^{i - m}}} \right.,{d_{i} \in \left\{ {0,1} \right\}}}{\left. {3.\mspace{14mu}{\sum\limits_{i = 0}^{l}{c_{i}/\left( 2^{w} \right)^{i}}}}\leftarrow{\sum\limits_{i = 0}^{m}{d_{i}/2^{i}}} \right.,\;{c_{i} \in \left\{ {0,1,\ldots\mspace{11mu},{2^{w} - 1}} \right\}}}{{{\left. {4.\mspace{14mu} D_{i}}\leftarrow{{\frac{i_{w - 1}}{2^{w - 1}}D} +} \right....} + {\frac{i_{0}}{2^{0}}D}},{{{for}\mspace{14mu} i} = {\left( {i_{w - 1}\mspace{14mu}\ldots\mspace{14mu} i_{0}} \right)_{2} \in \left\{ {{0,1},\ldots\mspace{11mu},{2^{w} - 1}} \right\}}}}\left. {5.\mspace{14mu} Q}\leftarrow{O6.\mspace{14mu}{for}\mspace{14mu} i\mspace{14mu}{from}\mspace{14mu} l\mspace{14mu}{downto}\mspace{14mu} 0\mspace{14mu}{do}\text{:}\;{for}\mspace{14mu} j\mspace{14mu}{from}\mspace{14mu} 1\mspace{14mu}{to}\mspace{14mu} w\mspace{14mu}{do}\text{:}\mspace{14mu} Q}\leftarrow{{HEC\_ HLV}(Q)Q}\leftarrow{Q + {D_{c_{i}}7.\mspace{14mu}{return}\mspace{14mu} Q}} \right.} & \left\lbrack {{Formula}\mspace{14mu} 41} \right\rbrack \end{matrix}$

Further, the inverse of divisor D=(U,V), U=x ² +u ₁ x+u ₀, and V=v ₁ x and v ₀, can be represented as follows: −D=(U,V+h mod U)=(U,(v ₁ +h ₂ u ₁ +h ₁)x+(v ₀ +h ₂ u ₀ +h ₀)). In particular, if h₂=1, no finite field multiplication is required, and four finite field addition operations suffices to find −D from D. The subtraction of the divisor D can be computed by the addition of the divisor −D. That is, the addition and subtraction of a divisor can be found with the same complexity.

Accordingly, it is possible to express a scalar value by also using negative values, and carry out scalar multiplication using the negative values. First, using NAK (Non-Adjacent Form), a given integer s is expressed by {−1, 0, 1}. In NAF, the scalar value of the integer s expressed in binary representation is scanned from the least significant bits. If there is a spot where number 1's appear adjacent to each other, this is expressed as follows, for example:

in the case of (11), this is expressed as (10-1), that is, 3=2²−1; and

in the case of (111), this is expressed as (100-1), that is, 7=2³−1.

The computation method for NAF is represented below as Algorithm 15a [Algorithm 15a].

$\begin{matrix} {{{Algorithm}\mspace{14mu} 15\; a\mspace{14mu}{Conversion}\mspace{14mu}{to}\mspace{14mu}{NAF}}{{{{Input}\text{:}\mspace{14mu} s} = {{\sum\limits_{j = 0}^{l - 1}{s_{j}2^{j}}} \in Z}},{s_{j} \in \left\{ {0,1} \right\}}}{{{{Output}\text{:}\mspace{14mu}{{NAF}(s)}} = {\sum\limits_{j = 0}^{l}{{s^{\prime}}_{j}2^{j}}}},{s_{j}^{\prime} \in \left\{ {{- 1},0,1} \right\}}}{\left. {1.\mspace{14mu} c_{0}}\leftarrow 0 \right.,{\left. s_{l}\leftarrow{0{2.\mspace{14mu}{For}\mspace{14mu} j}} \right. = {\left. {0\mspace{14mu}{to}\mspace{14mu} l\mspace{14mu}{do}\text{:}c_{j + 1}}\leftarrow{\left\lfloor {\left( {s_{j} + s_{j + 1} + c_{j}} \right)/2} \right\rfloor s_{j}^{\prime}}\leftarrow{s_{j} + c_{j} - {2\; c_{j + 1}3.\mspace{14mu}{return}\mspace{14mu}{{NAF}(S)}}} \right. = \left( {s_{l}^{\prime}s_{l - 1}^{\prime}\mspace{14mu}\ldots\mspace{14mu} s_{0}^{\prime}} \right)}}}} & \left\lbrack {{Formula}\mspace{14mu} 42} \right\rbrack \end{matrix}$

NAF represents an expression with the least number of non-zero bits. Since divisor addition or subtraction is performed at the portions of non-zero bits, scalar multiplication can be computed faster as the number of non-zero bits becomes smaller. The scalar value expression using NAF can be applied to the halve-and-add binary method and the halve-and-add window method. HEC_HLV used in each of the halve-and-add binary method and halve-and-add window method may be HEC_HLV of Algorithm 5a [Algorithm 5a] using a random curve, HEC_HLV with constraints h₂=1, f₄=0 provided to the curve parameters of Algorithm 8a [Algorithm 8a], HEC_HLV with constraints h₂=h₁=1, f₄=0 provided to the curve parameters of Algorithm 10a [Algorithm 10a], or HEC_HLV with constraints h₂=h₁=h₀=1, f₄=0 provided to the curve parameters of Algorithm 10a [Algorithm 10a]. The halve-and-add binary method using NAF is represented below as Algorithm 16a [Algorithm 16a].

$\begin{matrix} {{{{Algorithm}\mspace{14mu} 16\; a\mspace{14mu}{Halve}\text{-}{and}\text{-}{add}\mspace{14mu}{NAF}\mspace{14mu}{Binary}\mspace{14mu}{Method}}{{{Input}\text{:}\mspace{14mu} D} \in {{{J\left( F_{2^{n}} \right)}\mspace{14mu}{such}\mspace{14mu}{that}\mspace{14mu} 2\; D} \neq O}},{d \in Z},{r:\;{{order}\mspace{14mu}{of}\mspace{14mu} D}},{m = \left\lfloor {\log_{\; 2}r} \right\rfloor}}{{Output}\text{:}\mspace{14mu}{scalar}\mspace{14mu}{multiplication}\mspace{14mu}{dD}}\;{\left. {1.\mspace{14mu}{\sum\limits_{i = 0}^{m}{{\hat{d}}_{i}2^{i\;}}}}\leftarrow{{NAF}\left( {2^{m}d\;{mod}\; r} \right)} \right.,{{\hat{d}}_{i} \in \left\{ {{- 1},0,1} \right\}}}{\left. {2.\mspace{14mu}{\sum\limits_{i = 0}^{m}{d_{i}/2^{i}}}}\leftarrow{\sum\limits_{i = 0}^{m}{{\hat{d}}_{i}2^{i - m}}} \right.,{d_{i} \in \left\{ {{- 1},0,1} \right\}}}\left. {3.\mspace{14mu} Q}\leftarrow{O4.\mspace{14mu}{for}\mspace{14mu} i\mspace{14mu}{from}\mspace{14mu} m\mspace{14mu}{downto}\mspace{14mu} 0\mspace{14mu}{do}\text{:}\mspace{11mu} Q}\leftarrow{{HEC\_ HLV}(Q){{{if}\mspace{14mu} d_{i}} > {0\mspace{14mu}{then}\mspace{14mu} Q}}}\leftarrow{Q + {D{{{if}\mspace{14mu} d_{i}} < {0\mspace{14mu}{then}\mspace{14mu} Q}}}}\leftarrow{Q - {D5.\mspace{14mu}{return}\mspace{14mu} Q}} \right.} & \left\lbrack {{Formula}\mspace{14mu} 43} \right\rbrack \end{matrix}$

[Verification of Increased Computation Speed]

Next, the complexity of the computation applied to each of Processing Examples 7 to 11 described above is found, and verification is made as to an increase in computation speed.

In the case of HEC_HLV(h₂=1, f₄=0), the required complexity is, on average, 19.5M+2S+1I+3SR+2H+2T.

First, a case where a finite field is defined by a normal basis is considered. As described above, when using a normal basis, only the complexity of M and I may be taken into account. According to the document [A. Menezes. Elliptic Curve Public Key Cryptosystems. Kluwer Academic Publishers, 1993.], assuming that finite fields are Fq, q=2^(n), one inversion operation is equivalent to the number of multiplication operations computed by the following expression, that is: └ log₂(n−1)┘w(n−1)−1  [Formula 44]

In this case, w(n−1) denotes the number of 1's in the binary expression of n−1. For example, if n=83, 89, 113, then I=8M, and if n=103, then I=9M.

Here, assuming that I=8M, the complexity of HEC_HLV(h ₂=1, f ₄=0) is represented as 19.5M+1I=27.5M.

On the other hand, in the case of HarleyDBL, its complexity is represented as 21M+1I=29M, so HEC_HLV is about 5% faster than HarleyDBL. Further, when the table-lookup method is used, the complexity becomes 18M+1I=26M, so HEC_HLV is about 10% faster than HarleyDBL.

Further, in the case of HEC_HLV(h₂=h₁=1, f₄=0), the complexity is, on average, 14.5 M+3S+1I+2.5SR+2H+2T. In this case, 14.5M+1I=22.5M.

On the other hand, in the case of HarleyDBL, the complexity is represented as 18M+1I=26M, so HEC_HLV is about 13% faster than HarleyDBL. Further, when the table-lookup method is used, the complexity becomes 14M+1I=22M, so HEC_HLV is about 15% faster than HarleyDBL.

Further, the complexity of HEC_HLV(h₂=h₁=h₀=1, f₄=0) is, on average, 13.5M+3S+1I+2.5SR+2H+2T. In this case, 13.5M+1I=21.5M.

On the other hand, in the case of HarleyDBL, its complexity is 15M+1I=23M, so HEC_HLV is about 6% faster than HarleyDBL. Further, when the table-lookup method is used, the complexity becomes 17M+1I=25M, so HEC_HLV is about 14% faster than HarleyDBL.

Next, the complexity in the case of a polynomial basis will be evaluated. It is assumed that the complexities of S, I, SR, H, T are as follows: S=0.1M, I=8M, SR=0.5M, H=0.5M, T=0.5M. In the case of HEC_HLV(h₂=1, f₄=0), the complexity is, on average, 19.5M+2S+1I+2.5SR+2H+2T=29.95M.

On the other hand, in the case of HarleyDBL, the complexity is 21M+5S+1I=29.5M, so HarleyDBL is about 1% faster than HEC_HLV.

Further, when the table-lookup method is used, the complexity becomes 18M+2S+1I+2SR+2H=28.2M, so HEC_HLV is about 4% faster than HarleyDBL.

Further, in the case of HEC_HLV(h₂=h₁=1, f₄=0), the complexity is, on average, 14.5M+3S+1I+2.5SR+2H+2T=25.05M.

On the other hand, in the case of HarleyDBL, the complexity is 18M+7S+1I=26.7M, so HEC_HLV is about 6% faster than HarleyDBL.

Further, when the table-lookup method is used, the complexity becomes 14M+3S+1I+2SR+2H=24.3M, so HEC_HLV is about 9% faster than HarleyDBL.

Further, the complexity of HEC_HLV(h₂=h₁=h₀=1, f₄=0) is, on average, 13.5M+3S+1I+2.5SR+2H+2T=24.05M. On the other hand, in the case of HarleyDBL, its complexity is 15M+7S+1I=23.7M, so HarleyDBL is about 1% faster than HEC_HLV. Further, when the table-lookup method is used, the complexity becomes 13M+3S+1I+2SR+2H=23.3M, so HEC_HLV is about 2% faster than HarleyDBL.

Further, speed comparison was carried out for the case of a polynomial basis through software implementation.

The software implementation was carried out under the environment as indicated below:

CPU: PentiumII 300 MHx

OS: RedHat7.3

Compiler: gcc2.96.

The operations of M (multiplication) and S (squaring), I (inversion), SR (square root operation) and T (trace (determination as to whether roots exist for a quadratic equation)), and H (half-trace) (operation to find the root of a quadratic equation)) were carried out in the manner as disclosed in the following documents: [D. Hankerson, J. Hernandez, and A. Menezes. Software Implementation of Elliptic Curve Cryptography over Binary Fields. CHES 2000, LNCS 1965, pp. 1-24, 2000. Algorithm 4.6, 4.7]; [S. Shantz. From Euclid's GCD to Montgomery Multiplication to the Great Divide. TR-2001-95, Sun Microsystems, Inc., 2001.]; [K. Fong, D. Hankerson, J. Lopez, and A. Menezes. Field inversion operation and point halving revised. Technical Report CORR2003-18, www.cacr.math.uwaterloo.ca/techreports/2003/corr2003-18.pdf]; and [K. Fong, D. Hankerson, J. Lopez, and A. Menezes. Field inversion operation and point halving revised. Technical Report CORR2003-18, www.cacr.math.uwaterloo.ca/techreports/2003/corr2003-18.pdf Algorithm 4.7], respectively.

M, S, I, SR, H, T were implemented with respect to three finite fields of n=83, 89, 113, and the ratios to M were found. In this case, the following irreducible polynomials were used:

in the case of n=3, z ⁸³ +z ⁷ +z ⁴ +z ²+1=0;

in the case of n=9, z ⁸⁹ +z ³⁸+1=0; and

in the case of n=113, z ¹¹³ +z ⁹+1=0

The complexities in the respective cases were as follows.

n=83: S/M=0.12, I/M=7.96, SR/M=0.57, H/M=0.58

n=89: S/M=0.05, I/M=8.74, SR/M=0.14, H/M=0.61

n=113: S/M=0.06, I/M=8.56, SR/M=0.10, H/M=0.50

Applying these to the complexity of HarleyDBL, 21M+5S+1I, yields the following.

n=83: HarleyDBL 29.56M

n=89: HarleyDBL 29.99M

n=113: HarleyDBL 29.86M

Applying these to the complexity of HEC_HLV(h₂=1, f₄=0), 19.5M+2.5SR+2H+2T, yields the following.

n=83: HEC_HLV(h₂=1, f₄=0) 30.285M

n=89: HEC_HLV(h₂=1, f₄=0) 29.91M

n=113: HEC_HLV(h₂=1, f₄=0) 29.43M

In this case, when n=83, HarleyDBL is 2% faster than HEC_HLV. Further, when n-89, 113, HEC_HLV is about 0.3%, 1.5% faster than HarleyDBL, respectively.

Further, applying these to the complexity of HEC_HLV(h₂=1, f₄=0) to which the table-lookup method is applied, 18M+2S+1I+2SR+2H, yields the following.

n=83: HEC_HLV(h₂=1, f₄=0, with table-lookup) 28.5M

n=89: HEC_HLV(h₂=1, f₄=0, with table-lookup) 28.34M

n=113: HEC_HLV(h₂=1, f₄=0, with table-lookup) 27.88M

In this case, when n=83, 89, 113, HEC_HLV is 4%, 5%, 6% faster than Harley DBL, respectively.

Further, in the case of h₂=h₁=1, f₄=0, applying these to the complexity of HarleyDBL, 18M+7S+1I, yields the following.

n=83: HarleyDBL 27.4M

n=89: HarleyDBL 27.09M

n=113: HarleyDBL 26.98M

Next, applying these to the complexity of HEC_HLV(h₂=h₁=1, f₄=0), 14.5M+3S+1I+2.5SR+2H+2T, yields the following.

n=83: HEC_HLV(h₂=1, f₄=0) 25.405M

n=89: HEC_HLV(h₂=1, f₄=0) 24.96M

n=113: HEC_HLV(h₂=1, f₄=0) 24.49M

In this case, when n=83, 89, 113, HEC_HLV is 7%, 8%, 10% faster than HarleyDBL, respectively.

Further, applying these to the complexity of HEC_HLV(h₂=h₁=1, f₄=0) to which the table-lookup method is applied, 14M+3S+1I+2SR+2H, yields the following.

n=83: HEC_HLV(h₂=h₁=1, f₄=0, with table-lookup) 24.62M

n=89: HEC_HLV(h₂=h₁=1, f₄=0, with table-lookup) 24.39M

n=113: HEC_HLV(h₂=h₁=1, f₄=0, with table-lookup) 23.94M

In this case, when n=83, 89, 113, HEC_HLV is 10%, 8%, 11% faster than Harley DBL, respectively.

Further, in the case of h₂=h₁=h₀=1, f₄=0, applying these to the complexity of HarleyDBL, 15M+7S+1I, yields the following.

n=83: HarleyDBL 23.8M

n=89: HarleyDBL 24.09M

n=113: HarleyDBL 23.98M

Next, applying these to the complexity of HEC_HLV(h₂=h₁ h₀=1, f₄=0), 13.5M+3S+1I+2.5SR+2H+2T, yields the following.

n=83: HEC_HLV(h₂=h₁=h₀=1, f₄=0) 24.405M

n=89: HEC_HLV(h₂=h₁=h₀=1, f₄=0) 23.96M

n=113: HEC_HLV(h₂=h₁=h₀=1, f₄=0) 23.49M

In this case, when n=83, HarleyDBL is 2% faster than HEC_HLV. Further, when n=89, 113, HEC_HLV is 0.5%, 2% faster than Harley DBL, respectively.

Further, applying these to the complexity of HEC_HLV(h₂=h₁=h₀=1, f₄=0) to which the table-lookup method is applied, 13M+3S+1I+2SR+2H, yields the following.

n=83: HEC_HLV(h₂=h₁=h₀=1, f₄=0 with table-lookup) 23.62M

n=89: HEC_HLV(h₂=h₁=h₀=1, f₄=0 with table-lookup) 23.39M

n=113: HEC_HLV(h₂=h₁=h₀=1, f₄=0 with table-lookup) 22.94M

In this case, when n=83, 89, 113, HEC_HLV is 1%, 3%, 4% faster than Harley DBL, respectively.

From the foregoing, it can be said that HEC_HLV is faster than HarleyDBL in most of the cases. When the curve parameters are h₂=h₁=1, f₄=0, in particular, HEC_HLV is faster than HarleyDBL in all the cases.

Next, the complexity of scalar multiplication is considered. With regard to the above-mentioned examples, in the cases where HEC_HLV is faster than HarleyDBL, scalar multiplication using the combination of addition and halving is faster than scalar multiplication using the combination of addition and doubling. Now, a comparison will be made on the specific complexity of scalar multiplication in each individual case. As for the curve parameters, h₂=h₁=1, f4=0 are used. Further, as the scalar multiplication algorithm, the above-described NAF+binary method (Algorithm 16a [Algorithm 16a]) is used. Since the ratio of steps 1, 2 to the entire scalar multiplication process in this algorithm is very small, the complexity thereof is ignored. Here, the complexity is considered for the cases of n=83, 89, 113 for both a normal basis and a polynomial basis. Further, the order of the base point is assumed to be 165 bits, 177 bits, 225 bits with respect to n=83, 89, 113, respectively. Further, in the repeating portion of step 4, the repetition is made for the number of bits of the order of the base point. Divisor addition is carried out in the manner as disclosed in the document [T. Lange, Efficient arithmetic on genus 2 hyperelliptic curves over finite fields via explicit formulae. Cryptology ePrint Archive, 2002/121, IACR, 2002]. It should be noted that the curve parameters are h₂=h₁=1, f₄=0.

The complexity required for the divisor addition in this case is 21M+3S+1I. The scalar value is expressed by {−1, 0, 1} using NAF. If the scalar value is defined as m, there are about m/3 non-zero bits. Therefore, the complexity of NAF+binary method is computed as follows: ((the complexity of addition•subtraction)/3+(the complexity of halving or doubling))×(the number of bits of the order of the base point).

First, the case of a normal basis will be considered.

It is assumed that I=8M.

In the case of h₂=h₁=1, f₄=0,

n=83: addition•doubling: 5885M

n=89: addition•doubling: 6313M

n=113: addition•doubling: 8025M

In the case of h₂=h₁=1, f₄=0,

n=83: addition•halving: 5307.5M

n=89: addition•halving: 5693.5M

n=113: addition•halving: 7237.5M

In the case of h₂=h₁=1, f₄=0+table loop-up method,

n=83: addition•halving: 5225M

n=89: addition•halving: 5605M

n=113: addition•halving: 7125M

Next, the case of a polynomial basis is considered.

In the case of h₂=h₁=1, f₄=0,

n=83: addition•doubling: 6116M

n=89: addition•doubling: 6505.93M

n=113: addition•doubling: 8245.5M

In the case of h₂=h₁=1, f₄=0,

n=83: addition•halving: 5786.82M

n=89: addition•halving: 6128.92M

n=113: addition•halving: 7685.25M

In the case of h₂=h₁=1, f₄=0+table loop-up method

n=83: addition•halving: 5657.3M

n=89: addition•halving: 6028.03M

n=113: addition•halving: 7561.5M

It can be said that (scalar multiple of addition•halving) is faster than (scalar multiple of addition•doubling) by about 10 to 11% in the case of a normal basis, and by about 5 to 8% in the case of a polynomial basis.

As has been described above, according to the processing of the present invention, halving on elliptic curve cryptography is extended to hyperelliptic curve cryptography to thereby realize fast computation.

In the case of cryptographic computation employing computations on a divisor on a hyperelliptic curve, an arithmetic computation that puts a large load on the processing is the scalar multiplication of a divisor. In this regard, the processing according to the present invention as described above enables faster scalar multiplication to achieve a considerable improvement in the processing of hyperelliptic curve cryptography.

As described above, HECC (Hyper-Elliptic Curve Cryptography) is a generalized concept of ECC (Elliptic Curve Cryptography). Hence, the present invention can be applied to cryptographic processing using ECC (Elliptic Curve Cryptography) employed in a variety of applications, specifically including signature processing, generation of encrypted data, decryption, key pre-distribution system, and authentication processing. Faster computation can be achieved by replacing the portion of scalar multiplications in the computing processing of ECC (Elliptic Curve Cryptography) by the above-mentioned scalar multiplications.

[Functional Configuration of the Cryptographic System]

FIG. 7 is a block diagram showing the functional configuration of a cryptographic system according to the present invention. A cryptographic system 100 includes a base-point generating section 101 that generates a divisor D as a base point, a storage section 102 storing the table described with reference to Processing Example 5 mentioned above, that is, a table recording information as to which of k₁, k₁′, (k₀, k₀′) is correct on the basis of the computed value of [½^(i)D] with respect to a divisor D fixed in advance, and a computation executing section 103.

The computation executing section 103 executes computing operations including as computing processing in the computation of scalar multiplication with respect to a divisor D on a hyperelliptic curve. Specifically, the computation executing section executes computing operations including in the scalar multiplication with respect to a divisor on a hyperelliptic curve of genus 2 in characteristic 2 with random parameters. For example, computation executing section 103 executes computing operations including in the scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with parameters h(x)=x²+x+h₀, f₄=0, or in the scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 with a parameter h(x)=x.

[Applications of the Invention to an Algorithm for Generating and Verifying a Digital Signature]

The following description explains a case in which the computation technique according to the present invention is applied to the scalar multiplication of an ECDSA (EC-Digital Signature Algorithm), which is an algorithm for generation and verification of a digital signature and to which elliptic curve cryptography is applied, as an example of a specific cryptographic algorithm to which the processing of the present invention can be applied. In accordance with the IEEE1363, a digital signature is generated and verified by execution of the following sequence.

(1): Inputs

(1-1): Input domain parameters and a base point G (order r) of an elliptic curve

(1-2): Input a secret key s of the signatory.

(1-3): Input a plain text M.

(2): Generation of a Key

(2-1): Use W=sG as a public key for the secret key s.

(3): Generation of a Signature

(3-1): Generate a random integer u, where 0<u<r.

(3-2): Compute V=uG=(x_(v), y_(v)).

(3-3): Convert x_(v) into an integer i.

(3-4): Compute c=i mod r. If c=0, go to step 3-1.

(3-5): f=h (M), where h is a hash function.

(3-6): Compute d=u⁻¹ (f+sc) mod r. If d=0, go to step 3-1.

(3-7): Use (c, d) as a signature for the plain text M.

(4): Verification of a Signature

(4-1): Check whether relations 0<c<r and 0<d<r hold true. If they do not hold true, output “invalid”.

(4-2): Compute h=d⁻¹ mod r, h₁=fh mod r, and h₂=ch mod r.

(4-3): Compute P=(x_(p), y_(p))=h₁G+h₂W If P=0, output “invalid”.

(4-4): Convert x_(p) into an integer i.

(4-5): Compute c′=i mod r.

(4-6): If c′=c, output “valid”. Otherwise, output “invalid”.

The proposed techniques using a hyperelliptic curve can be applied to the following steps of the above algorithm:

(2-1): Use W=sG as a public key for the secret key s;

(3-2): Compute V=uG=(x_(v), y_(v)); and

(4-3): Compute P=(x_(p), y_(p))=h₁G+h₂W If P=0, output “invalid”.

The computing processing W=sG, V=uG, and P=(x_(p), y_(p))=h₁G+h₂W in each of the respective steps (2-1), (3-2), and (4-3) represents scalar multiplication processing on a divisor and can be carried out faster through the application of the present invention. Further, the computing processing sG, uG, and h₁G represents scalar multiplication processing on a fixed divisor and can be carried out faster through the application of the table-lookup method according to the present invention.

[Hardware Configuration Example of the Cryptographic System]

Finally, an example configuration of an IC module 200 serving as a device for executing the cryptographic processing described above will be described with reference to FIG. 8. The processing described above can be executed by a variety of information processing apparatus such as a PC, an IC card, and a reader/writer. The IC module 200 shown in FIG. 8 can be incorporated in these information-processing apparatus.

A CPU (Central Processing Unit) 201 shown in FIG. 8 is a processor for executing a variety of programs to start and end cryptographic processing, control transmissions and receptions of data, and control transfers of data among respective components. A memory 202 includes a ROM (Read Only Memory) and a RAM (Random Access Memory). The ROM stores programs to be executed by the CPU 201 or fixed data as computational parameters. The RAM is used as the storage area/work area for storing programs to be executed by the CPU 201 to carry out processing, and computational parameters that change as the processing of programs proceeds.

It should be noted that a computation executing program stored in the memory 202 is set as a program including the sequence of execution of the addition and doubling, which are performed as the base point setting processing and the scalar multiplication processing described above. In addition, the memory 202 can also serve as a storage area for key data or the like required for cryptographic processing. It is preferable to design the storage area for data or the like as a memory area having a tamper-proof structure.

A cryptographic section 203 executes processing such as cryptographic processing including the scalar multiplication described above and decryption processing. It should be noted that while the cryptographic section 203 is shown as an independent module, such an independent cryptographic processing module may not be provided. That is, a configuration may be adopted in which, for example, a cryptographic program can be stored in the ROM, and the CPU 201 reads out the cryptographic program from the ROM and executes the program.

A random-number generator 204 executes the processing of generating a random number required for the generation of a key or the like necessary for cryptographic processing.

A transmitting/receiving section 205 is a data-communication processing unit for executing data communications with an external apparatus. The transmitting/receiving section 205 executes data communications with an IC module such as a reader/writer, and executes such processing as the outputting of encrypted text generated in the IC module or inputting of data from an external apparatus such as a reader/writer.

The present invention has been described in detail so far with reference to its specific embodiments. It is obvious, however, that a person skilled in the art can anticipate various modifications and alternatives to the embodiments without departing from the scope of the present invention. That is, the foregoing disclosure of the present invention has been made only by way of examples and should not be construed restrictively. The scope of the present invention should be determined by reference to the appended claims.

The series of processing described in this specification can be executed by hardware, software or a composite configuration of both. If the series of processing is to be executed by software, the series of processing can be executed by installing a program recording the processing sequence into a memory of a computer built in dedicated hardware, or by installing the program into a general purpose computer capable of executing various processing.

For example, the program may be stored in advance in a hard disc or a ROM (Read Only Memory) as a recording medium. Alternatively, the program may be stored (recorded) temporarily or permanently in a removable recording medium such as a flexible disc, a CD-ROM (Compact Disc Read Only Memory), an MO (Magneto-optical) disc, a DVD (Digital Versatile Disc), a magnetic disc, and a semiconductor memory. Such a removable recording medium can be provided in the form of so-called package software.

Other than being installed into a computer from the above-described removable recording medium, the program can be wireless-transferred to a computer from a download site, or wired-transferred to a computer via a network such as a LAN (Local Area Network) or the Internet. The computer receives the program thus transferred and installs the program into a built-in recording medium such as a hard disk.

The various processing described in this specification may be executed not only time sequentially in the order as they appear in the description but may be executed in parallel or independently depending on the throughput of the device executing the processes. Further, the term system as used in this specification refers to a logical assembly of a plurality of devices, and is not limited to one in which devices of respective configurations are located within the same casing.

INDUSTRIAL APPLICABILITY

According to the configuration of the present invention, halving on elliptic curve cryptography is extended to hyperelliptic curve cryptography to thereby realize fast computation. In the case of cryptographic computation employing computations on a divisor on a hyperelliptic curve, a computing operation that puts a large load on the processing is the scalar multiplication of a divisor. Hence, by realizing faster scalar multiplication by the processing according to the present invention as described above, a considerable improvement can be achieved in terms of the processing of hyperelliptic curve cryptography. The present invention can be thus applied to apparatuses, devices, and the like required to perform fast and secure cryptographic computations, such as an IC card.

According to the configuration of the present invention, in scalar multiplication with respect to a divisor D in hyperelliptic curve cryptography, faster scalar multiplication can be realized by executing computing operations including halving as computing processing. For example, fast computation is realized by executing computing operations including halving in scalar multiplication with respect to a divisor D on a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x²+x+h₀, f₄=0 as parameters, a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x²+h₁x+h₀, f₄=0 as parameters, or a hyperelliptic curve of genus 2 in characteristic 2 having h(x)=x as a parameter. The present invention can be thus applied to apparatuses, devices, and the like required to perform fast and secure cryptographic computations, such as an IC card.

According to the configuration of the present invention, a further reduction in the complexity of scalar multiplication of a divisor and hence faster computation can be achieved through the application of a table that records which of k₁, k₁′, (k₀, k₀′) is correct on the basis of a computed value of [½^(i)D] with respect to a divisor D fixed in advance. The present invention can be thus applied to apparatuses, devices, and the like required to perform fast and secure cryptographic computations, such as an IC card.

According to the configuration of the present invention, in scalar multiplication with respect to a divisor D in hyperelliptic curve cryptography, computing operations including halving are executed as computing processing, and an algorithm for reducing the number of inversion operations executed in the halving computation processing is applied, thereby making it possible to achieve a further reduction in the complexity of scalar multiplication of a divisor and hence faster computation.

It should be understood that various changes and modifications to the presently preferred embodiments described herein will be apparent to those skilled in the art. Such changes and modifications can be made without departing from the spirit and scope of the present invention and without diminishing its intended advantages. It is therefore intended that such changes and modifications be covered by the appended claims. 

1. A method of operating a cryptographic system including instructions, the method comprising: (a) causing a base point generator to generate a divisor of a hyperelliptic curve of genus 2 in characteristic 2, the hyperelliptic curve being defined by y²+h(x)y=f(x), the divisor being equal to a formal sum of points, wherein: (a) y is a first point on the hyperelliptic curve; (b) x is a second point on the hyperelliptic curve; (c) h(x) is a first polynomial; and (d) f(x) is a second polynomial; (b) causing a processor to execute the instructions to cause the processor to perform halving on the generated divisor of the hyperelliptic curve of genus 2 in characteristic 2; and (c) causing the processor to execute the instructions to cause the processor to perform scalar multiplication using the halved divisor of the hyperelliptic curve of genus 2 in characteristic
 2. 2. The method of claim 1, wherein the first polynomial includes a random integer.
 3. The method of claim 1, wherein the h(x)=x²+x+h₀.
 4. The method of claim 1, wherein the h(x)=x²+h₁x+h₀.
 5. The method of claim 1, wherein the h(x)=x.
 6. The method of claim 1, which includes causing the processor to execute the instructions to cause the processor to reduce the complexity of the halving by determining processing based on a lookup of a table that records which of k₁, k₁′, (k₀, k₀′) is correct based on a computed value of [½^(i)D] with respect to a divisor D fixed in advance.
 7. The method of claim 1, which includes causing the processor to execute the instructions to cause the processor to calculate a value of an inverse 1/k₁ by multiplication and addition processing without performing an inversion, by application of the following relational expression: 1/k ₁ =h ₂ +k ₁ u ₂₁, which is derived from a halving computation algorithm in which Input: D₂=(U₂, V₂), and Output: D₁=(U₁, V₁)=[½]D₂, where U_(i)(x)=x²+u_(i1)x+u_(i0), Vi(x)=v_(i1)x+v_(i0), gcd(h, U_(i))=1, i=1,
 2. 8. The method of claim 1, which includes causing the processor to execute the instructions to cause the processor to execute a computation according to an algorithm having a setting for not applying 1/u₂₁ as an input value, in a halving computation algorithm in which Input: D₂=(U₂, V₂); and Output: D₁=(U₁, V₁)=[½]D₂, where U_(i)(x)=x²+u_(i1)x+u_(i0), Vi(x)=v_(i1)x+v_(i0), gcd(h, U_(i))=1, i=1,
 2. 9. The method of claim 1, which includes causing the processor to execute the instructions to cause the processor to: (a) set, as an input value, 1/h₁ ² that is a previously calculated value; and (b) apply the previously calculated input value 1/h₁ ² without executing processing of calculating an inverse 1/h₁ ², wherein the h(x)=x²+h₁x+h₀.
 10. A cryptographic system comprising: a processor; a base point generator operatively connected to the processor, the base point generator which generates a divisor of a hyperelliptic curve of genus 2 in characteristic 2, the hyperelliptic curve being defined by y²+h(x)y=f(x), the divisor being equal to a formal sum of points, wherein: (a) y is a first point on the hyperelliptic curve; (b) x is a second point on the hyperelliptic curve; (c) h(x) is a first polynomial; and (d) f(x) is a second polynomial; and a memory device operatively coupled to the processor, the memory device storing instructions which when executed by the processor, cause the processor, in cooperation with the memory device to: (a) perform halving on the generated divisor of the hyperelliptic curve of genus 2 in characteristic 2; and (b) perform scalar multiplication using the halved divisor of the hyperelliptic curve of genus 2 in characteristic
 2. 11. The cryptographic system of claim 10, wherein the first polynomial includes a random integer.
 12. The cryptographic system of claim 10, wherein the h(x)=x²+x+h₀.
 13. The cryptographic system of claim 10, wherein the h(x)=x²+h₁x+h₀.
 14. The cryptographic system of claim 10, wherein the h(x)=x.
 15. The cryptographic system of claim 10, wherein the instructions, when executed by the processor, cause the processor to reduce the complexity of the halving by determining processing based on a lookup of a table recording which of k₁, k₁′, (k₀, k₀′) is correct based on a computed value of [½^(i)D] with respect to a divisor D fixed in advance.
 16. The cryptographic system of claim 10, wherein the instructions, when executed by the processor, cause the processor to calculate a value of an inverse 1/k₁ by multiplication and addition processing without performing an inversion, by application of the following relational expression: 1/k ₁ =h ₂ +k ₁ u ₂₁, which is derived from a halving computation algorithm in which Input: D₂=(U₂, V₂), and Output: D₁=(U₁, V₁)=[½]D₂, where U_(i)(x)=x²+u_(i1)x+u_(i0), Vi(x)=v_(i1)x+v_(i0), gcd(h, U_(i))=1, i=1,
 2. 17. The cryptographic system of claim 10, wherein the instructions, when executed by the processor, cause the processor to execute: (a) a halving computation algorithm in which Input: D₂=(U₂, V₂), and Output: D₁=(U₁, V₁)=[½]D₂, where U_(i)(x)=x²+u_(i1)x+u_(i0), Vi(x)=v_(i1)x+v_(i0), gcd(h, U₁)=1, i=1, 2; and (b) computation according to an algorithm having a setting for not applying 1/u₂₁ as an input value.
 18. The cryptographic system of claim 10, wherein: (a) h(x)=x²+h₁x+h₀; and (b) the instructions, when executed by the processor, cause the processor to execute computation to which, with 1/h₁ ² that is a previously calculated value being set as an input value, the previously calculated input value 1/h₁ ² is applied without executing processing of calculating an inverse 1/h₁ ².
 19. A non-transitory computer-readable medium storing instructions structured to cause a computer to: (a) generate a divisor of a hyperelliptic curve of genus 2 in characteristic 2, the hyperelliptic curve being defined by y²+h(x)y=f(x), the divisor being equal to a formal sum of points, wherein: (a) y is a first point on the hyperelliptic curve; (b) x is a second point on the hyperelliptic curve; (c) h(x) is a first polynomial; and (d) f(x) is a second polynomial; and (b) perform halving on the generated divisor of the hyperelliptic curve of genus 2 in characteristic 2; and (c) perform scalar multiplication using the halved divisor of the hyperelliptic curve of genus 2 in characteristic
 2. 